diff options
Diffstat (limited to 'Documentation/vm')
| -rw-r--r-- | Documentation/vm/00-INDEX | 36 | ||||
| -rw-r--r-- | Documentation/vm/active_mm.txt | 83 | ||||
| -rw-r--r-- | Documentation/vm/balance | 93 | ||||
| -rw-r--r-- | Documentation/vm/cleancache.txt | 279 | ||||
| -rw-r--r-- | Documentation/vm/highmem.txt | 162 | ||||
| -rw-r--r-- | Documentation/vm/hugetlbpage.txt | 309 | ||||
| -rw-r--r-- | Documentation/vm/hwpoison.txt | 182 | ||||
| -rw-r--r-- | Documentation/vm/ksm.txt | 82 | ||||
| -rw-r--r-- | Documentation/vm/locking | 130 | ||||
| -rw-r--r-- | Documentation/vm/numa | 149 | ||||
| -rw-r--r-- | Documentation/vm/numa_memory_policy.txt | 453 | ||||
| -rw-r--r-- | Documentation/vm/overcommit-accounting | 73 | ||||
| -rw-r--r-- | Documentation/vm/page_migration | 149 | ||||
| -rw-r--r-- | Documentation/vm/pagemap.txt | 151 | ||||
| -rw-r--r-- | Documentation/vm/slub.txt | 283 | ||||
| -rw-r--r-- | Documentation/vm/transhuge.txt | 299 | ||||
| -rw-r--r-- | Documentation/vm/unevictable-lru.txt | 690 |
17 files changed, 3603 insertions, 0 deletions
diff --git a/Documentation/vm/00-INDEX b/Documentation/vm/00-INDEX new file mode 100644 index 00000000..5481c8ba --- /dev/null +++ b/Documentation/vm/00-INDEX @@ -0,0 +1,36 @@ +00-INDEX + - this file. +active_mm.txt + - An explanation from Linus about tsk->active_mm vs tsk->mm. +balance + - various information on memory balancing. +hugepage-mmap.c + - Example app using huge page memory with the mmap system call. +hugepage-shm.c + - Example app using huge page memory with Sys V shared memory system calls. +hugetlbpage.txt + - a brief summary of hugetlbpage support in the Linux kernel. +hwpoison.txt + - explains what hwpoison is +ksm.txt + - how to use the Kernel Samepage Merging feature. +locking + - info on how locking and synchronization is done in the Linux vm code. +map_hugetlb.c + - an example program that uses the MAP_HUGETLB mmap flag. +numa + - information about NUMA specific code in the Linux vm. +numa_memory_policy.txt + - documentation of concepts and APIs of the 2.6 memory policy support. +overcommit-accounting + - description of the Linux kernels overcommit handling modes. +page-types.c + - Tool for querying page flags +page_migration + - description of page migration in NUMA systems. +pagemap.txt + - pagemap, from the userspace perspective +slub.txt + - a short users guide for SLUB. +unevictable-lru.txt + - Unevictable LRU infrastructure diff --git a/Documentation/vm/active_mm.txt b/Documentation/vm/active_mm.txt new file mode 100644 index 00000000..dbf45817 --- /dev/null +++ b/Documentation/vm/active_mm.txt @@ -0,0 +1,83 @@ +List: linux-kernel +Subject: Re: active_mm +From: Linus Torvalds <torvalds () transmeta ! com> +Date: 1999-07-30 21:36:24 + +Cc'd to linux-kernel, because I don't write explanations all that often, +and when I do I feel better about more people reading them. + +On Fri, 30 Jul 1999, David Mosberger wrote: +> +> Is there a brief description someplace on how "mm" vs. "active_mm" in +> the task_struct are supposed to be used? (My apologies if this was +> discussed on the mailing lists---I just returned from vacation and +> wasn't able to follow linux-kernel for a while). + +Basically, the new setup is: + + - we have "real address spaces" and "anonymous address spaces". The + difference is that an anonymous address space doesn't care about the + user-level page tables at all, so when we do a context switch into an + anonymous address space we just leave the previous address space + active. + + The obvious use for a "anonymous address space" is any thread that + doesn't need any user mappings - all kernel threads basically fall into + this category, but even "real" threads can temporarily say that for + some amount of time they are not going to be interested in user space, + and that the scheduler might as well try to avoid wasting time on + switching the VM state around. Currently only the old-style bdflush + sync does that. + + - "tsk->mm" points to the "real address space". For an anonymous process, + tsk->mm will be NULL, for the logical reason that an anonymous process + really doesn't _have_ a real address space at all. + + - however, we obviously need to keep track of which address space we + "stole" for such an anonymous user. For that, we have "tsk->active_mm", + which shows what the currently active address space is. + + The rule is that for a process with a real address space (ie tsk->mm is + non-NULL) the active_mm obviously always has to be the same as the real + one. + + For a anonymous process, tsk->mm == NULL, and tsk->active_mm is the + "borrowed" mm while the anonymous process is running. When the + anonymous process gets scheduled away, the borrowed address space is + returned and cleared. + +To support all that, the "struct mm_struct" now has two counters: a +"mm_users" counter that is how many "real address space users" there are, +and a "mm_count" counter that is the number of "lazy" users (ie anonymous +users) plus one if there are any real users. + +Usually there is at least one real user, but it could be that the real +user exited on another CPU while a lazy user was still active, so you do +actually get cases where you have a address space that is _only_ used by +lazy users. That is often a short-lived state, because once that thread +gets scheduled away in favour of a real thread, the "zombie" mm gets +released because "mm_users" becomes zero. + +Also, a new rule is that _nobody_ ever has "init_mm" as a real MM any +more. "init_mm" should be considered just a "lazy context when no other +context is available", and in fact it is mainly used just at bootup when +no real VM has yet been created. So code that used to check + + if (current->mm == &init_mm) + +should generally just do + + if (!current->mm) + +instead (which makes more sense anyway - the test is basically one of "do +we have a user context", and is generally done by the page fault handler +and things like that). + +Anyway, I put a pre-patch-2.3.13-1 on ftp.kernel.org just a moment ago, +because it slightly changes the interfaces to accommodate the alpha (who +would have thought it, but the alpha actually ends up having one of the +ugliest context switch codes - unlike the other architectures where the MM +and register state is separate, the alpha PALcode joins the two, and you +need to switch both together). + +(From http://marc.info/?l=linux-kernel&m=93337278602211&w=2) diff --git a/Documentation/vm/balance b/Documentation/vm/balance new file mode 100644 index 00000000..c46e68cf --- /dev/null +++ b/Documentation/vm/balance @@ -0,0 +1,93 @@ +Started Jan 2000 by Kanoj Sarcar <kanoj@sgi.com> + +Memory balancing is needed for non __GFP_WAIT as well as for non +__GFP_IO allocations. + +There are two reasons to be requesting non __GFP_WAIT allocations: +the caller can not sleep (typically intr context), or does not want +to incur cost overheads of page stealing and possible swap io for +whatever reasons. + +__GFP_IO allocation requests are made to prevent file system deadlocks. + +In the absence of non sleepable allocation requests, it seems detrimental +to be doing balancing. Page reclamation can be kicked off lazily, that +is, only when needed (aka zone free memory is 0), instead of making it +a proactive process. + +That being said, the kernel should try to fulfill requests for direct +mapped pages from the direct mapped pool, instead of falling back on +the dma pool, so as to keep the dma pool filled for dma requests (atomic +or not). A similar argument applies to highmem and direct mapped pages. +OTOH, if there is a lot of free dma pages, it is preferable to satisfy +regular memory requests by allocating one from the dma pool, instead +of incurring the overhead of regular zone balancing. + +In 2.2, memory balancing/page reclamation would kick off only when the +_total_ number of free pages fell below 1/64 th of total memory. With the +right ratio of dma and regular memory, it is quite possible that balancing +would not be done even when the dma zone was completely empty. 2.2 has +been running production machines of varying memory sizes, and seems to be +doing fine even with the presence of this problem. In 2.3, due to +HIGHMEM, this problem is aggravated. + +In 2.3, zone balancing can be done in one of two ways: depending on the +zone size (and possibly of the size of lower class zones), we can decide +at init time how many free pages we should aim for while balancing any +zone. The good part is, while balancing, we do not need to look at sizes +of lower class zones, the bad part is, we might do too frequent balancing +due to ignoring possibly lower usage in the lower class zones. Also, +with a slight change in the allocation routine, it is possible to reduce +the memclass() macro to be a simple equality. + +Another possible solution is that we balance only when the free memory +of a zone _and_ all its lower class zones falls below 1/64th of the +total memory in the zone and its lower class zones. This fixes the 2.2 +balancing problem, and stays as close to 2.2 behavior as possible. Also, +the balancing algorithm works the same way on the various architectures, +which have different numbers and types of zones. If we wanted to get +fancy, we could assign different weights to free pages in different +zones in the future. + +Note that if the size of the regular zone is huge compared to dma zone, +it becomes less significant to consider the free dma pages while +deciding whether to balance the regular zone. The first solution +becomes more attractive then. + +The appended patch implements the second solution. It also "fixes" two +problems: first, kswapd is woken up as in 2.2 on low memory conditions +for non-sleepable allocations. Second, the HIGHMEM zone is also balanced, +so as to give a fighting chance for replace_with_highmem() to get a +HIGHMEM page, as well as to ensure that HIGHMEM allocations do not +fall back into regular zone. This also makes sure that HIGHMEM pages +are not leaked (for example, in situations where a HIGHMEM page is in +the swapcache but is not being used by anyone) + +kswapd also needs to know about the zones it should balance. kswapd is +primarily needed in a situation where balancing can not be done, +probably because all allocation requests are coming from intr context +and all process contexts are sleeping. For 2.3, kswapd does not really +need to balance the highmem zone, since intr context does not request +highmem pages. kswapd looks at the zone_wake_kswapd field in the zone +structure to decide whether a zone needs balancing. + +Page stealing from process memory and shm is done if stealing the page would +alleviate memory pressure on any zone in the page's node that has fallen below +its watermark. + +watemark[WMARK_MIN/WMARK_LOW/WMARK_HIGH]/low_on_memory/zone_wake_kswapd: These +are per-zone fields, used to determine when a zone needs to be balanced. When +the number of pages falls below watermark[WMARK_MIN], the hysteric field +low_on_memory gets set. This stays set till the number of free pages becomes +watermark[WMARK_HIGH]. When low_on_memory is set, page allocation requests will +try to free some pages in the zone (providing GFP_WAIT is set in the request). +Orthogonal to this, is the decision to poke kswapd to free some zone pages. +That decision is not hysteresis based, and is done when the number of free +pages is below watermark[WMARK_LOW]; in which case zone_wake_kswapd is also set. + + +(Good) Ideas that I have heard: +1. Dynamic experience should influence balancing: number of failed requests +for a zone can be tracked and fed into the balancing scheme (jalvo@mbay.net) +2. Implement a replace_with_highmem()-like replace_with_regular() to preserve +dma pages. (lkd@tantalophile.demon.co.uk) diff --git a/Documentation/vm/cleancache.txt b/Documentation/vm/cleancache.txt new file mode 100644 index 00000000..142fbb0f --- /dev/null +++ b/Documentation/vm/cleancache.txt @@ -0,0 +1,279 @@ +MOTIVATION + +Cleancache is a new optional feature provided by the VFS layer that +potentially dramatically increases page cache effectiveness for +many workloads in many environments at a negligible cost. + +Cleancache can be thought of as a page-granularity victim cache for clean +pages that the kernel's pageframe replacement algorithm (PFRA) would like +to keep around, but can't since there isn't enough memory. So when the +PFRA "evicts" a page, it first attempts to use cleancache code to +put the data contained in that page into "transcendent memory", memory +that is not directly accessible or addressable by the kernel and is +of unknown and possibly time-varying size. + +Later, when a cleancache-enabled filesystem wishes to access a page +in a file on disk, it first checks cleancache to see if it already +contains it; if it does, the page of data is copied into the kernel +and a disk access is avoided. + +Transcendent memory "drivers" for cleancache are currently implemented +in Xen (using hypervisor memory) and zcache (using in-kernel compressed +memory) and other implementations are in development. + +FAQs are included below. + +IMPLEMENTATION OVERVIEW + +A cleancache "backend" that provides transcendent memory registers itself +to the kernel's cleancache "frontend" by calling cleancache_register_ops, +passing a pointer to a cleancache_ops structure with funcs set appropriately. +Note that cleancache_register_ops returns the previous settings so that +chaining can be performed if desired. The functions provided must conform to +certain semantics as follows: + +Most important, cleancache is "ephemeral". Pages which are copied into +cleancache have an indefinite lifetime which is completely unknowable +by the kernel and so may or may not still be in cleancache at any later time. +Thus, as its name implies, cleancache is not suitable for dirty pages. +Cleancache has complete discretion over what pages to preserve and what +pages to discard and when. + +Mounting a cleancache-enabled filesystem should call "init_fs" to obtain a +pool id which, if positive, must be saved in the filesystem's superblock; +a negative return value indicates failure. A "put_page" will copy a +(presumably about-to-be-evicted) page into cleancache and associate it with +the pool id, a file key, and a page index into the file. (The combination +of a pool id, a file key, and an index is sometimes called a "handle".) +A "get_page" will copy the page, if found, from cleancache into kernel memory. +An "invalidate_page" will ensure the page no longer is present in cleancache; +an "invalidate_inode" will invalidate all pages associated with the specified +file; and, when a filesystem is unmounted, an "invalidate_fs" will invalidate +all pages in all files specified by the given pool id and also surrender +the pool id. + +An "init_shared_fs", like init_fs, obtains a pool id but tells cleancache +to treat the pool as shared using a 128-bit UUID as a key. On systems +that may run multiple kernels (such as hard partitioned or virtualized +systems) that may share a clustered filesystem, and where cleancache +may be shared among those kernels, calls to init_shared_fs that specify the +same UUID will receive the same pool id, thus allowing the pages to +be shared. Note that any security requirements must be imposed outside +of the kernel (e.g. by "tools" that control cleancache). Or a +cleancache implementation can simply disable shared_init by always +returning a negative value. + +If a get_page is successful on a non-shared pool, the page is invalidated +(thus making cleancache an "exclusive" cache). On a shared pool, the page +is NOT invalidated on a successful get_page so that it remains accessible to +other sharers. The kernel is responsible for ensuring coherency between +cleancache (shared or not), the page cache, and the filesystem, using +cleancache invalidate operations as required. + +Note that cleancache must enforce put-put-get coherency and get-get +coherency. For the former, if two puts are made to the same handle but +with different data, say AAA by the first put and BBB by the second, a +subsequent get can never return the stale data (AAA). For get-get coherency, +if a get for a given handle fails, subsequent gets for that handle will +never succeed unless preceded by a successful put with that handle. + +Last, cleancache provides no SMP serialization guarantees; if two +different Linux threads are simultaneously putting and invalidating a page +with the same handle, the results are indeterminate. Callers must +lock the page to ensure serial behavior. + +CLEANCACHE PERFORMANCE METRICS + +If properly configured, monitoring of cleancache is done via debugfs in +the /sys/kernel/debug/mm/cleancache directory. The effectiveness of cleancache +can be measured (across all filesystems) with: + +succ_gets - number of gets that were successful +failed_gets - number of gets that failed +puts - number of puts attempted (all "succeed") +invalidates - number of invalidates attempted + +A backend implementation may provide additional metrics. + +FAQ + +1) Where's the value? (Andrew Morton) + +Cleancache provides a significant performance benefit to many workloads +in many environments with negligible overhead by improving the +effectiveness of the pagecache. Clean pagecache pages are +saved in transcendent memory (RAM that is otherwise not directly +addressable to the kernel); fetching those pages later avoids "refaults" +and thus disk reads. + +Cleancache (and its sister code "frontswap") provide interfaces for +this transcendent memory (aka "tmem"), which conceptually lies between +fast kernel-directly-addressable RAM and slower DMA/asynchronous devices. +Disallowing direct kernel or userland reads/writes to tmem +is ideal when data is transformed to a different form and size (such +as with compression) or secretly moved (as might be useful for write- +balancing for some RAM-like devices). Evicted page-cache pages (and +swap pages) are a great use for this kind of slower-than-RAM-but-much- +faster-than-disk transcendent memory, and the cleancache (and frontswap) +"page-object-oriented" specification provides a nice way to read and +write -- and indirectly "name" -- the pages. + +In the virtual case, the whole point of virtualization is to statistically +multiplex physical resources across the varying demands of multiple +virtual machines. This is really hard to do with RAM and efforts to +do it well with no kernel change have essentially failed (except in some +well-publicized special-case workloads). Cleancache -- and frontswap -- +with a fairly small impact on the kernel, provide a huge amount +of flexibility for more dynamic, flexible RAM multiplexing. +Specifically, the Xen Transcendent Memory backend allows otherwise +"fallow" hypervisor-owned RAM to not only be "time-shared" between multiple +virtual machines, but the pages can be compressed and deduplicated to +optimize RAM utilization. And when guest OS's are induced to surrender +underutilized RAM (e.g. with "self-ballooning"), page cache pages +are the first to go, and cleancache allows those pages to be +saved and reclaimed if overall host system memory conditions allow. + +And the identical interface used for cleancache can be used in +physical systems as well. The zcache driver acts as a memory-hungry +device that stores pages of data in a compressed state. And +the proposed "RAMster" driver shares RAM across multiple physical +systems. + +2) Why does cleancache have its sticky fingers so deep inside the + filesystems and VFS? (Andrew Morton and Christoph Hellwig) + +The core hooks for cleancache in VFS are in most cases a single line +and the minimum set are placed precisely where needed to maintain +coherency (via cleancache_invalidate operations) between cleancache, +the page cache, and disk. All hooks compile into nothingness if +cleancache is config'ed off and turn into a function-pointer- +compare-to-NULL if config'ed on but no backend claims the ops +functions, or to a compare-struct-element-to-negative if a +backend claims the ops functions but a filesystem doesn't enable +cleancache. + +Some filesystems are built entirely on top of VFS and the hooks +in VFS are sufficient, so don't require an "init_fs" hook; the +initial implementation of cleancache didn't provide this hook. +But for some filesystems (such as btrfs), the VFS hooks are +incomplete and one or more hooks in fs-specific code are required. +And for some other filesystems, such as tmpfs, cleancache may +be counterproductive. So it seemed prudent to require a filesystem +to "opt in" to use cleancache, which requires adding a hook in +each filesystem. Not all filesystems are supported by cleancache +only because they haven't been tested. The existing set should +be sufficient to validate the concept, the opt-in approach means +that untested filesystems are not affected, and the hooks in the +existing filesystems should make it very easy to add more +filesystems in the future. + +The total impact of the hooks to existing fs and mm files is only +about 40 lines added (not counting comments and blank lines). + +3) Why not make cleancache asynchronous and batched so it can + more easily interface with real devices with DMA instead + of copying each individual page? (Minchan Kim) + +The one-page-at-a-time copy semantics simplifies the implementation +on both the frontend and backend and also allows the backend to +do fancy things on-the-fly like page compression and +page deduplication. And since the data is "gone" (copied into/out +of the pageframe) before the cleancache get/put call returns, +a great deal of race conditions and potential coherency issues +are avoided. While the interface seems odd for a "real device" +or for real kernel-addressable RAM, it makes perfect sense for +transcendent memory. + +4) Why is non-shared cleancache "exclusive"? And where is the + page "invalidated" after a "get"? (Minchan Kim) + +The main reason is to free up space in transcendent memory and +to avoid unnecessary cleancache_invalidate calls. If you want inclusive, +the page can be "put" immediately following the "get". If +put-after-get for inclusive becomes common, the interface could +be easily extended to add a "get_no_invalidate" call. + +The invalidate is done by the cleancache backend implementation. + +5) What's the performance impact? + +Performance analysis has been presented at OLS'09 and LCA'10. +Briefly, performance gains can be significant on most workloads, +especially when memory pressure is high (e.g. when RAM is +overcommitted in a virtual workload); and because the hooks are +invoked primarily in place of or in addition to a disk read/write, +overhead is negligible even in worst case workloads. Basically +cleancache replaces I/O with memory-copy-CPU-overhead; on older +single-core systems with slow memory-copy speeds, cleancache +has little value, but in newer multicore machines, especially +consolidated/virtualized machines, it has great value. + +6) How do I add cleancache support for filesystem X? (Boaz Harrash) + +Filesystems that are well-behaved and conform to certain +restrictions can utilize cleancache simply by making a call to +cleancache_init_fs at mount time. Unusual, misbehaving, or +poorly layered filesystems must either add additional hooks +and/or undergo extensive additional testing... or should just +not enable the optional cleancache. + +Some points for a filesystem to consider: + +- The FS should be block-device-based (e.g. a ram-based FS such + as tmpfs should not enable cleancache) +- To ensure coherency/correctness, the FS must ensure that all + file removal or truncation operations either go through VFS or + add hooks to do the equivalent cleancache "invalidate" operations +- To ensure coherency/correctness, either inode numbers must + be unique across the lifetime of the on-disk file OR the + FS must provide an "encode_fh" function. +- The FS must call the VFS superblock alloc and deactivate routines + or add hooks to do the equivalent cleancache calls done there. +- To maximize performance, all pages fetched from the FS should + go through the do_mpag_readpage routine or the FS should add + hooks to do the equivalent (cf. btrfs) +- Currently, the FS blocksize must be the same as PAGESIZE. This + is not an architectural restriction, but no backends currently + support anything different. +- A clustered FS should invoke the "shared_init_fs" cleancache + hook to get best performance for some backends. + +7) Why not use the KVA of the inode as the key? (Christoph Hellwig) + +If cleancache would use the inode virtual address instead of +inode/filehandle, the pool id could be eliminated. But, this +won't work because cleancache retains pagecache data pages +persistently even when the inode has been pruned from the +inode unused list, and only invalidates the data page if the file +gets removed/truncated. So if cleancache used the inode kva, +there would be potential coherency issues if/when the inode +kva is reused for a different file. Alternately, if cleancache +invalidated the pages when the inode kva was freed, much of the value +of cleancache would be lost because the cache of pages in cleanache +is potentially much larger than the kernel pagecache and is most +useful if the pages survive inode cache removal. + +8) Why is a global variable required? + +The cleancache_enabled flag is checked in all of the frequently-used +cleancache hooks. The alternative is a function call to check a static +variable. Since cleancache is enabled dynamically at runtime, systems +that don't enable cleancache would suffer thousands (possibly +tens-of-thousands) of unnecessary function calls per second. So the +global variable allows cleancache to be enabled by default at compile +time, but have insignificant performance impact when cleancache remains +disabled at runtime. + +9) Does cleanache work with KVM? + +The memory model of KVM is sufficiently different that a cleancache +backend may have less value for KVM. This remains to be tested, +especially in an overcommitted system. + +10) Does cleancache work in userspace? It sounds useful for + memory hungry caches like web browsers. (Jamie Lokier) + +No plans yet, though we agree it sounds useful, at least for +apps that bypass the page cache (e.g. O_DIRECT). + +Last updated: Dan Magenheimer, April 13 2011 diff --git a/Documentation/vm/highmem.txt b/Documentation/vm/highmem.txt new file mode 100644 index 00000000..4324d24f --- /dev/null +++ b/Documentation/vm/highmem.txt @@ -0,0 +1,162 @@ + + ==================== + HIGH MEMORY HANDLING + ==================== + +By: Peter Zijlstra <a.p.zijlstra@chello.nl> + +Contents: + + (*) What is high memory? + + (*) Temporary virtual mappings. + + (*) Using kmap_atomic. + + (*) Cost of temporary mappings. + + (*) i386 PAE. + + +==================== +WHAT IS HIGH MEMORY? +==================== + +High memory (highmem) is used when the size of physical memory approaches or +exceeds the maximum size of virtual memory. At that point it becomes +impossible for the kernel to keep all of the available physical memory mapped +at all times. This means the kernel needs to start using temporary mappings of +the pieces of physical memory that it wants to access. + +The part of (physical) memory not covered by a permanent mapping is what we +refer to as 'highmem'. There are various architecture dependent constraints on +where exactly that border lies. + +In the i386 arch, for example, we choose to map the kernel into every process's +VM space so that we don't have to pay the full TLB invalidation costs for +kernel entry/exit. This means the available virtual memory space (4GiB on +i386) has to be divided between user and kernel space. + +The traditional split for architectures using this approach is 3:1, 3GiB for +userspace and the top 1GiB for kernel space: + + +--------+ 0xffffffff + | Kernel | + +--------+ 0xc0000000 + | | + | User | + | | + +--------+ 0x00000000 + +This means that the kernel can at most map 1GiB of physical memory at any one +time, but because we need virtual address space for other things - including +temporary maps to access the rest of the physical memory - the actual direct +map will typically be less (usually around ~896MiB). + +Other architectures that have mm context tagged TLBs can have separate kernel +and user maps. Some hardware (like some ARMs), however, have limited virtual +space when they use mm context tags. + + +========================== +TEMPORARY VIRTUAL MAPPINGS +========================== + +The kernel contains several ways of creating temporary mappings: + + (*) vmap(). This can be used to make a long duration mapping of multiple + physical pages into a contiguous virtual space. It needs global + synchronization to unmap. + + (*) kmap(). This permits a short duration mapping of a single page. It needs + global synchronization, but is amortized somewhat. It is also prone to + deadlocks when using in a nested fashion, and so it is not recommended for + new code. + + (*) kmap_atomic(). This permits a very short duration mapping of a single + page. Since the mapping is restricted to the CPU that issued it, it + performs well, but the issuing task is therefore required to stay on that + CPU until it has finished, lest some other task displace its mappings. + + kmap_atomic() may also be used by interrupt contexts, since it is does not + sleep and the caller may not sleep until after kunmap_atomic() is called. + + It may be assumed that k[un]map_atomic() won't fail. + + +================= +USING KMAP_ATOMIC +================= + +When and where to use kmap_atomic() is straightforward. It is used when code +wants to access the contents of a page that might be allocated from high memory +(see __GFP_HIGHMEM), for example a page in the pagecache. The API has two +functions, and they can be used in a manner similar to the following: + + /* Find the page of interest. */ + struct page *page = find_get_page(mapping, offset); + + /* Gain access to the contents of that page. */ + void *vaddr = kmap_atomic(page); + + /* Do something to the contents of that page. */ + memset(vaddr, 0, PAGE_SIZE); + + /* Unmap that page. */ + kunmap_atomic(vaddr); + +Note that the kunmap_atomic() call takes the result of the kmap_atomic() call +not the argument. + +If you need to map two pages because you want to copy from one page to +another you need to keep the kmap_atomic calls strictly nested, like: + + vaddr1 = kmap_atomic(page1); + vaddr2 = kmap_atomic(page2); + + memcpy(vaddr1, vaddr2, PAGE_SIZE); + + kunmap_atomic(vaddr2); + kunmap_atomic(vaddr1); + + +========================== +COST OF TEMPORARY MAPPINGS +========================== + +The cost of creating temporary mappings can be quite high. The arch has to +manipulate the kernel's page tables, the data TLB and/or the MMU's registers. + +If CONFIG_HIGHMEM is not set, then the kernel will try and create a mapping +simply with a bit of arithmetic that will convert the page struct address into +a pointer to the page contents rather than juggling mappings about. In such a +case, the unmap operation may be a null operation. + +If CONFIG_MMU is not set, then there can be no temporary mappings and no +highmem. In such a case, the arithmetic approach will also be used. + + +======== +i386 PAE +======== + +The i386 arch, under some circumstances, will permit you to stick up to 64GiB +of RAM into your 32-bit machine. This has a number of consequences: + + (*) Linux needs a page-frame structure for each page in the system and the + pageframes need to live in the permanent mapping, which means: + + (*) you can have 896M/sizeof(struct page) page-frames at most; with struct + page being 32-bytes that would end up being something in the order of 112G + worth of pages; the kernel, however, needs to store more than just + page-frames in that memory... + + (*) PAE makes your page tables larger - which slows the system down as more + data has to be accessed to traverse in TLB fills and the like. One + advantage is that PAE has more PTE bits and can provide advanced features + like NX and PAT. + +The general recommendation is that you don't use more than 8GiB on a 32-bit +machine - although more might work for you and your workload, you're pretty +much on your own - don't expect kernel developers to really care much if things +come apart. diff --git a/Documentation/vm/hugetlbpage.txt b/Documentation/vm/hugetlbpage.txt new file mode 100644 index 00000000..f8551b38 --- /dev/null +++ b/Documentation/vm/hugetlbpage.txt @@ -0,0 +1,309 @@ + +The intent of this file is to give a brief summary of hugetlbpage support in +the Linux kernel. This support is built on top of multiple page size support +that is provided by most modern architectures. For example, i386 +architecture supports 4K and 4M (2M in PAE mode) page sizes, ia64 +architecture supports multiple page sizes 4K, 8K, 64K, 256K, 1M, 4M, 16M, +256M and ppc64 supports 4K and 16M. A TLB is a cache of virtual-to-physical +translations. Typically this is a very scarce resource on processor. +Operating systems try to make best use of limited number of TLB resources. +This optimization is more critical now as bigger and bigger physical memories +(several GBs) are more readily available. + +Users can use the huge page support in Linux kernel by either using the mmap +system call or standard SYSV shared memory system calls (shmget, shmat). + +First the Linux kernel needs to be built with the CONFIG_HUGETLBFS +(present under "File systems") and CONFIG_HUGETLB_PAGE (selected +automatically when CONFIG_HUGETLBFS is selected) configuration +options. + +The /proc/meminfo file provides information about the total number of +persistent hugetlb pages in the kernel's huge page pool. It also displays +information about the number of free, reserved and surplus huge pages and the +default huge page size. The huge page size is needed for generating the +proper alignment and size of the arguments to system calls that map huge page +regions. + +The output of "cat /proc/meminfo" will include lines like: + +..... +HugePages_Total: vvv +HugePages_Free: www +HugePages_Rsvd: xxx +HugePages_Surp: yyy +Hugepagesize: zzz kB + +where: +HugePages_Total is the size of the pool of huge pages. +HugePages_Free is the number of huge pages in the pool that are not yet + allocated. +HugePages_Rsvd is short for "reserved," and is the number of huge pages for + which a commitment to allocate from the pool has been made, + but no allocation has yet been made. Reserved huge pages + guarantee that an application will be able to allocate a + huge page from the pool of huge pages at fault time. +HugePages_Surp is short for "surplus," and is the number of huge pages in + the pool above the value in /proc/sys/vm/nr_hugepages. The + maximum number of surplus huge pages is controlled by + /proc/sys/vm/nr_overcommit_hugepages. + +/proc/filesystems should also show a filesystem of type "hugetlbfs" configured +in the kernel. + +/proc/sys/vm/nr_hugepages indicates the current number of "persistent" huge +pages in the kernel's huge page pool. "Persistent" huge pages will be +returned to the huge page pool when freed by a task. A user with root +privileges can dynamically allocate more or free some persistent huge pages +by increasing or decreasing the value of 'nr_hugepages'. + +Pages that are used as huge pages are reserved inside the kernel and cannot +be used for other purposes. Huge pages cannot be swapped out under +memory pressure. + +Once a number of huge pages have been pre-allocated to the kernel huge page +pool, a user with appropriate privilege can use either the mmap system call +or shared memory system calls to use the huge pages. See the discussion of +Using Huge Pages, below. + +The administrator can allocate persistent huge pages on the kernel boot +command line by specifying the "hugepages=N" parameter, where 'N' = the +number of huge pages requested. This is the most reliable method of +allocating huge pages as memory has not yet become fragmented. + +Some platforms support multiple huge page sizes. To allocate huge pages +of a specific size, one must precede the huge pages boot command parameters +with a huge page size selection parameter "hugepagesz=<size>". <size> must +be specified in bytes with optional scale suffix [kKmMgG]. The default huge +page size may be selected with the "default_hugepagesz=<size>" boot parameter. + +When multiple huge page sizes are supported, /proc/sys/vm/nr_hugepages +indicates the current number of pre-allocated huge pages of the default size. +Thus, one can use the following command to dynamically allocate/deallocate +default sized persistent huge pages: + + echo 20 > /proc/sys/vm/nr_hugepages + +This command will try to adjust the number of default sized huge pages in the +huge page pool to 20, allocating or freeing huge pages, as required. + +On a NUMA platform, the kernel will attempt to distribute the huge page pool +over all the set of allowed nodes specified by the NUMA memory policy of the +task that modifies nr_hugepages. The default for the allowed nodes--when the +task has default memory policy--is all on-line nodes with memory. Allowed +nodes with insufficient available, contiguous memory for a huge page will be +silently skipped when allocating persistent huge pages. See the discussion +below of the interaction of task memory policy, cpusets and per node attributes +with the allocation and freeing of persistent huge pages. + +The success or failure of huge page allocation depends on the amount of +physically contiguous memory that is present in system at the time of the +allocation attempt. If the kernel is unable to allocate huge pages from +some nodes in a NUMA system, it will attempt to make up the difference by +allocating extra pages on other nodes with sufficient available contiguous +memory, if any. + +System administrators may want to put this command in one of the local rc +init files. This will enable the kernel to allocate huge pages early in +the boot process when the possibility of getting physical contiguous pages +is still very high. Administrators can verify the number of huge pages +actually allocated by checking the sysctl or meminfo. To check the per node +distribution of huge pages in a NUMA system, use: + + cat /sys/devices/system/node/node*/meminfo | fgrep Huge + +/proc/sys/vm/nr_overcommit_hugepages specifies how large the pool of +huge pages can grow, if more huge pages than /proc/sys/vm/nr_hugepages are +requested by applications. Writing any non-zero value into this file +indicates that the hugetlb subsystem is allowed to try to obtain that +number of "surplus" huge pages from the kernel's normal page pool, when the +persistent huge page pool is exhausted. As these surplus huge pages become +unused, they are freed back to the kernel's normal page pool. + +When increasing the huge page pool size via nr_hugepages, any existing surplus +pages will first be promoted to persistent huge pages. Then, additional +huge pages will be allocated, if necessary and if possible, to fulfill +the new persistent huge page pool size. + +The administrator may shrink the pool of persistent huge pages for +the default huge page size by setting the nr_hugepages sysctl to a +smaller value. The kernel will attempt to balance the freeing of huge pages +across all nodes in the memory policy of the task modifying nr_hugepages. +Any free huge pages on the selected nodes will be freed back to the kernel's +normal page pool. + +Caveat: Shrinking the persistent huge page pool via nr_hugepages such that +it becomes less than the number of huge pages in use will convert the balance +of the in-use huge pages to surplus huge pages. This will occur even if +the number of surplus pages it would exceed the overcommit value. As long as +this condition holds--that is, until nr_hugepages+nr_overcommit_hugepages is +increased sufficiently, or the surplus huge pages go out of use and are freed-- +no more surplus huge pages will be allowed to be allocated. + +With support for multiple huge page pools at run-time available, much of +the huge page userspace interface in /proc/sys/vm has been duplicated in sysfs. +The /proc interfaces discussed above have been retained for backwards +compatibility. The root huge page control directory in sysfs is: + + /sys/kernel/mm/hugepages + +For each huge page size supported by the running kernel, a subdirectory +will exist, of the form: + + hugepages-${size}kB + +Inside each of these directories, the same set of files will exist: + + nr_hugepages + nr_hugepages_mempolicy + nr_overcommit_hugepages + free_hugepages + resv_hugepages + surplus_hugepages + +which function as described above for the default huge page-sized case. + + +Interaction of Task Memory Policy with Huge Page Allocation/Freeing + +Whether huge pages are allocated and freed via the /proc interface or +the /sysfs interface using the nr_hugepages_mempolicy attribute, the NUMA +nodes from which huge pages are allocated or freed are controlled by the +NUMA memory policy of the task that modifies the nr_hugepages_mempolicy +sysctl or attribute. When the nr_hugepages attribute is used, mempolicy +is ignored. + +The recommended method to allocate or free huge pages to/from the kernel +huge page pool, using the nr_hugepages example above, is: + + numactl --interleave <node-list> echo 20 \ + >/proc/sys/vm/nr_hugepages_mempolicy + +or, more succinctly: + + numactl -m <node-list> echo 20 >/proc/sys/vm/nr_hugepages_mempolicy + +This will allocate or free abs(20 - nr_hugepages) to or from the nodes +specified in <node-list>, depending on whether number of persistent huge pages +is initially less than or greater than 20, respectively. No huge pages will be +allocated nor freed on any node not included in the specified <node-list>. + +When adjusting the persistent hugepage count via nr_hugepages_mempolicy, any +memory policy mode--bind, preferred, local or interleave--may be used. The +resulting effect on persistent huge page allocation is as follows: + +1) Regardless of mempolicy mode [see Documentation/vm/numa_memory_policy.txt], + persistent huge pages will be distributed across the node or nodes + specified in the mempolicy as if "interleave" had been specified. + However, if a node in the policy does not contain sufficient contiguous + memory for a huge page, the allocation will not "fallback" to the nearest + neighbor node with sufficient contiguous memory. To do this would cause + undesirable imbalance in the distribution of the huge page pool, or + possibly, allocation of persistent huge pages on nodes not allowed by + the task's memory policy. + +2) One or more nodes may be specified with the bind or interleave policy. + If more than one node is specified with the preferred policy, only the + lowest numeric id will be used. Local policy will select the node where + the task is running at the time the nodes_allowed mask is constructed. + For local policy to be deterministic, the task must be bound to a cpu or + cpus in a single node. Otherwise, the task could be migrated to some + other node at any time after launch and the resulting node will be + indeterminate. Thus, local policy is not very useful for this purpose. + Any of the other mempolicy modes may be used to specify a single node. + +3) The nodes allowed mask will be derived from any non-default task mempolicy, + whether this policy was set explicitly by the task itself or one of its + ancestors, such as numactl. This means that if the task is invoked from a + shell with non-default policy, that policy will be used. One can specify a + node list of "all" with numactl --interleave or --membind [-m] to achieve + interleaving over all nodes in the system or cpuset. + +4) Any task mempolicy specifed--e.g., using numactl--will be constrained by + the resource limits of any cpuset in which the task runs. Thus, there will + be no way for a task with non-default policy running in a cpuset with a + subset of the system nodes to allocate huge pages outside the cpuset + without first moving to a cpuset that contains all of the desired nodes. + +5) Boot-time huge page allocation attempts to distribute the requested number + of huge pages over all on-lines nodes with memory. + +Per Node Hugepages Attributes + +A subset of the contents of the root huge page control directory in sysfs, +described above, will be replicated under each the system device of each +NUMA node with memory in: + + /sys/devices/system/node/node[0-9]*/hugepages/ + +Under this directory, the subdirectory for each supported huge page size +contains the following attribute files: + + nr_hugepages + free_hugepages + surplus_hugepages + +The free_' and surplus_' attribute files are read-only. They return the number +of free and surplus [overcommitted] huge pages, respectively, on the parent +node. + +The nr_hugepages attribute returns the total number of huge pages on the +specified node. When this attribute is written, the number of persistent huge +pages on the parent node will be adjusted to the specified value, if sufficient +resources exist, regardless of the task's mempolicy or cpuset constraints. + +Note that the number of overcommit and reserve pages remain global quantities, +as we don't know until fault time, when the faulting task's mempolicy is +applied, from which node the huge page allocation will be attempted. + + +Using Huge Pages + +If the user applications are going to request huge pages using mmap system +call, then it is required that system administrator mount a file system of +type hugetlbfs: + + mount -t hugetlbfs \ + -o uid=<value>,gid=<value>,mode=<value>,size=<value>,nr_inodes=<value> \ + none /mnt/huge + +This command mounts a (pseudo) filesystem of type hugetlbfs on the directory +/mnt/huge. Any files created on /mnt/huge uses huge pages. The uid and gid +options sets the owner and group of the root of the file system. By default +the uid and gid of the current process are taken. The mode option sets the +mode of root of file system to value & 0777. This value is given in octal. +By default the value 0755 is picked. The size option sets the maximum value of +memory (huge pages) allowed for that filesystem (/mnt/huge). The size is +rounded down to HPAGE_SIZE. The option nr_inodes sets the maximum number of +inodes that /mnt/huge can use. If the size or nr_inodes option is not +provided on command line then no limits are set. For size and nr_inodes +options, you can use [G|g]/[M|m]/[K|k] to represent giga/mega/kilo. For +example, size=2K has the same meaning as size=2048. + +While read system calls are supported on files that reside on hugetlb +file systems, write system calls are not. + +Regular chown, chgrp, and chmod commands (with right permissions) could be +used to change the file attributes on hugetlbfs. + +Also, it is important to note that no such mount command is required if the +applications are going to use only shmat/shmget system calls or mmap with +MAP_HUGETLB. Users who wish to use hugetlb page via shared memory segment +should be a member of a supplementary group and system admin needs to +configure that gid into /proc/sys/vm/hugetlb_shm_group. It is possible for +same or different applications to use any combination of mmaps and shm* +calls, though the mount of filesystem will be required for using mmap calls +without MAP_HUGETLB. For an example of how to use mmap with MAP_HUGETLB see +map_hugetlb.c. + +******************************************************************* + +/* + * hugepage-shm: see Documentation/vm/hugepage-shm.c + */ + +******************************************************************* + +/* + * hugepage-mmap: see Documentation/vm/hugepage-mmap.c + */ diff --git a/Documentation/vm/hwpoison.txt b/Documentation/vm/hwpoison.txt new file mode 100644 index 00000000..55006846 --- /dev/null +++ b/Documentation/vm/hwpoison.txt @@ -0,0 +1,182 @@ +What is hwpoison? + +Upcoming Intel CPUs have support for recovering from some memory errors +(``MCA recovery''). This requires the OS to declare a page "poisoned", +kill the processes associated with it and avoid using it in the future. + +This patchkit implements the necessary infrastructure in the VM. + +To quote the overview comment: + + * High level machine check handler. Handles pages reported by the + * hardware as being corrupted usually due to a 2bit ECC memory or cache + * failure. + * + * This focusses on pages detected as corrupted in the background. + * When the current CPU tries to consume corruption the currently + * running process can just be killed directly instead. This implies + * that if the error cannot be handled for some reason it's safe to + * just ignore it because no corruption has been consumed yet. Instead + * when that happens another machine check will happen. + * + * Handles page cache pages in various states. The tricky part + * here is that we can access any page asynchronous to other VM + * users, because memory failures could happen anytime and anywhere, + * possibly violating some of their assumptions. This is why this code + * has to be extremely careful. Generally it tries to use normal locking + * rules, as in get the standard locks, even if that means the + * error handling takes potentially a long time. + * + * Some of the operations here are somewhat inefficient and have non + * linear algorithmic complexity, because the data structures have not + * been optimized for this case. This is in particular the case + * for the mapping from a vma to a process. Since this case is expected + * to be rare we hope we can get away with this. + +The code consists of a the high level handler in mm/memory-failure.c, +a new page poison bit and various checks in the VM to handle poisoned +pages. + +The main target right now is KVM guests, but it works for all kinds +of applications. KVM support requires a recent qemu-kvm release. + +For the KVM use there was need for a new signal type so that +KVM can inject the machine check into the guest with the proper +address. This in theory allows other applications to handle +memory failures too. The expection is that near all applications +won't do that, but some very specialized ones might. + +--- + +There are two (actually three) modi memory failure recovery can be in: + +vm.memory_failure_recovery sysctl set to zero: + All memory failures cause a panic. Do not attempt recovery. + (on x86 this can be also affected by the tolerant level of the + MCE subsystem) + +early kill + (can be controlled globally and per process) + Send SIGBUS to the application as soon as the error is detected + This allows applications who can process memory errors in a gentle + way (e.g. drop affected object) + This is the mode used by KVM qemu. + +late kill + Send SIGBUS when the application runs into the corrupted page. + This is best for memory error unaware applications and default + Note some pages are always handled as late kill. + +--- + +User control: + +vm.memory_failure_recovery + See sysctl.txt + +vm.memory_failure_early_kill + Enable early kill mode globally + +PR_MCE_KILL + Set early/late kill mode/revert to system default + arg1: PR_MCE_KILL_CLEAR: Revert to system default + arg1: PR_MCE_KILL_SET: arg2 defines thread specific mode + PR_MCE_KILL_EARLY: Early kill + PR_MCE_KILL_LATE: Late kill + PR_MCE_KILL_DEFAULT: Use system global default +PR_MCE_KILL_GET + return current mode + + +--- + +Testing: + +madvise(MADV_HWPOISON, ....) + (as root) + Poison a page in the process for testing + + +hwpoison-inject module through debugfs + +/sys/debug/hwpoison/ + +corrupt-pfn + +Inject hwpoison fault at PFN echoed into this file. This does +some early filtering to avoid corrupted unintended pages in test suites. + +unpoison-pfn + +Software-unpoison page at PFN echoed into this file. This +way a page can be reused again. +This only works for Linux injected failures, not for real +memory failures. + +Note these injection interfaces are not stable and might change between +kernel versions + +corrupt-filter-dev-major +corrupt-filter-dev-minor + +Only handle memory failures to pages associated with the file system defined +by block device major/minor. -1U is the wildcard value. +This should be only used for testing with artificial injection. + +corrupt-filter-memcg + +Limit injection to pages owned by memgroup. Specified by inode number +of the memcg. + +Example: + mkdir /sys/fs/cgroup/mem/hwpoison + + usemem -m 100 -s 1000 & + echo `jobs -p` > /sys/fs/cgroup/mem/hwpoison/tasks + + memcg_ino=$(ls -id /sys/fs/cgroup/mem/hwpoison | cut -f1 -d' ') + echo $memcg_ino > /debug/hwpoison/corrupt-filter-memcg + + page-types -p `pidof init` --hwpoison # shall do nothing + page-types -p `pidof usemem` --hwpoison # poison its pages + +corrupt-filter-flags-mask +corrupt-filter-flags-value + +When specified, only poison pages if ((page_flags & mask) == value). +This allows stress testing of many kinds of pages. The page_flags +are the same as in /proc/kpageflags. The flag bits are defined in +include/linux/kernel-page-flags.h and documented in +Documentation/vm/pagemap.txt + +Architecture specific MCE injector + +x86 has mce-inject, mce-test + +Some portable hwpoison test programs in mce-test, see blow. + +--- + +References: + +http://halobates.de/mce-lc09-2.pdf + Overview presentation from LinuxCon 09 + +git://git.kernel.org/pub/scm/utils/cpu/mce/mce-test.git + Test suite (hwpoison specific portable tests in tsrc) + +git://git.kernel.org/pub/scm/utils/cpu/mce/mce-inject.git + x86 specific injector + + +--- + +Limitations: + +- Not all page types are supported and never will. Most kernel internal +objects cannot be recovered, only LRU pages for now. +- Right now hugepage support is missing. + +--- +Andi Kleen, Oct 2009 + diff --git a/Documentation/vm/ksm.txt b/Documentation/vm/ksm.txt new file mode 100644 index 00000000..b392e496 --- /dev/null +++ b/Documentation/vm/ksm.txt @@ -0,0 +1,82 @@ +How to use the Kernel Samepage Merging feature +---------------------------------------------- + +KSM is a memory-saving de-duplication feature, enabled by CONFIG_KSM=y, +added to the Linux kernel in 2.6.32. See mm/ksm.c for its implementation, +and http://lwn.net/Articles/306704/ and http://lwn.net/Articles/330589/ + +The KSM daemon ksmd periodically scans those areas of user memory which +have been registered with it, looking for pages of identical content which +can be replaced by a single write-protected page (which is automatically +copied if a process later wants to update its content). + +KSM was originally developed for use with KVM (where it was known as +Kernel Shared Memory), to fit more virtual machines into physical memory, +by sharing the data common between them. But it can be useful to any +application which generates many instances of the same data. + +KSM only merges anonymous (private) pages, never pagecache (file) pages. +KSM's merged pages were originally locked into kernel memory, but can now +be swapped out just like other user pages (but sharing is broken when they +are swapped back in: ksmd must rediscover their identity and merge again). + +KSM only operates on those areas of address space which an application +has advised to be likely candidates for merging, by using the madvise(2) +system call: int madvise(addr, length, MADV_MERGEABLE). + +The app may call int madvise(addr, length, MADV_UNMERGEABLE) to cancel +that advice and restore unshared pages: whereupon KSM unmerges whatever +it merged in that range. Note: this unmerging call may suddenly require +more memory than is available - possibly failing with EAGAIN, but more +probably arousing the Out-Of-Memory killer. + +If KSM is not configured into the running kernel, madvise MADV_MERGEABLE +and MADV_UNMERGEABLE simply fail with EINVAL. If the running kernel was +built with CONFIG_KSM=y, those calls will normally succeed: even if the +the KSM daemon is not currently running, MADV_MERGEABLE still registers +the range for whenever the KSM daemon is started; even if the range +cannot contain any pages which KSM could actually merge; even if +MADV_UNMERGEABLE is applied to a range which was never MADV_MERGEABLE. + +Like other madvise calls, they are intended for use on mapped areas of +the user address space: they will report ENOMEM if the specified range +includes unmapped gaps (though working on the intervening mapped areas), +and might fail with EAGAIN if not enough memory for internal structures. + +Applications should be considerate in their use of MADV_MERGEABLE, +restricting its use to areas likely to benefit. KSM's scans may use a lot +of processing power: some installations will disable KSM for that reason. + +The KSM daemon is controlled by sysfs files in /sys/kernel/mm/ksm/, +readable by all but writable only by root: + +pages_to_scan - how many present pages to scan before ksmd goes to sleep + e.g. "echo 100 > /sys/kernel/mm/ksm/pages_to_scan" + Default: 100 (chosen for demonstration purposes) + +sleep_millisecs - how many milliseconds ksmd should sleep before next scan + e.g. "echo 20 > /sys/kernel/mm/ksm/sleep_millisecs" + Default: 20 (chosen for demonstration purposes) + +run - set 0 to stop ksmd from running but keep merged pages, + set 1 to run ksmd e.g. "echo 1 > /sys/kernel/mm/ksm/run", + set 2 to stop ksmd and unmerge all pages currently merged, + but leave mergeable areas registered for next run + Default: 0 (must be changed to 1 to activate KSM, + except if CONFIG_SYSFS is disabled) + +The effectiveness of KSM and MADV_MERGEABLE is shown in /sys/kernel/mm/ksm/: + +pages_shared - how many shared pages are being used +pages_sharing - how many more sites are sharing them i.e. how much saved +pages_unshared - how many pages unique but repeatedly checked for merging +pages_volatile - how many pages changing too fast to be placed in a tree +full_scans - how many times all mergeable areas have been scanned + +A high ratio of pages_sharing to pages_shared indicates good sharing, but +a high ratio of pages_unshared to pages_sharing indicates wasted effort. +pages_volatile embraces several different kinds of activity, but a high +proportion there would also indicate poor use of madvise MADV_MERGEABLE. + +Izik Eidus, +Hugh Dickins, 17 Nov 2009 diff --git a/Documentation/vm/locking b/Documentation/vm/locking new file mode 100644 index 00000000..f61228bd --- /dev/null +++ b/Documentation/vm/locking @@ -0,0 +1,130 @@ +Started Oct 1999 by Kanoj Sarcar <kanojsarcar@yahoo.com> + +The intent of this file is to have an uptodate, running commentary +from different people about how locking and synchronization is done +in the Linux vm code. + +page_table_lock & mmap_sem +-------------------------------------- + +Page stealers pick processes out of the process pool and scan for +the best process to steal pages from. To guarantee the existence +of the victim mm, a mm_count inc and a mmdrop are done in swap_out(). +Page stealers hold kernel_lock to protect against a bunch of races. +The vma list of the victim mm is also scanned by the stealer, +and the page_table_lock is used to preserve list sanity against the +process adding/deleting to the list. This also guarantees existence +of the vma. Vma existence is not guaranteed once try_to_swap_out() +drops the page_table_lock. To guarantee the existence of the underlying +file structure, a get_file is done before the swapout() method is +invoked. The page passed into swapout() is guaranteed not to be reused +for a different purpose because the page reference count due to being +present in the user's pte is not released till after swapout() returns. + +Any code that modifies the vmlist, or the vm_start/vm_end/ +vm_flags:VM_LOCKED/vm_next of any vma *in the list* must prevent +kswapd from looking at the chain. + +The rules are: +1. To scan the vmlist (look but don't touch) you must hold the + mmap_sem with read bias, i.e. down_read(&mm->mmap_sem) +2. To modify the vmlist you need to hold the mmap_sem with + read&write bias, i.e. down_write(&mm->mmap_sem) *AND* + you need to take the page_table_lock. +3. The swapper takes _just_ the page_table_lock, this is done + because the mmap_sem can be an extremely long lived lock + and the swapper just cannot sleep on that. +4. The exception to this rule is expand_stack, which just + takes the read lock and the page_table_lock, this is ok + because it doesn't really modify fields anybody relies on. +5. You must be able to guarantee that while holding page_table_lock + or page_table_lock of mm A, you will not try to get either lock + for mm B. + +The caveats are: +1. find_vma() makes use of, and updates, the mmap_cache pointer hint. +The update of mmap_cache is racy (page stealer can race with other code +that invokes find_vma with mmap_sem held), but that is okay, since it +is a hint. This can be fixed, if desired, by having find_vma grab the +page_table_lock. + + +Code that add/delete elements from the vmlist chain are +1. callers of insert_vm_struct +2. callers of merge_segments +3. callers of avl_remove + +Code that changes vm_start/vm_end/vm_flags:VM_LOCKED of vma's on +the list: +1. expand_stack +2. mprotect +3. mlock +4. mremap + +It is advisable that changes to vm_start/vm_end be protected, although +in some cases it is not really needed. Eg, vm_start is modified by +expand_stack(), it is hard to come up with a destructive scenario without +having the vmlist protection in this case. + +The page_table_lock nests with the inode i_mmap_mutex and the kmem cache +c_spinlock spinlocks. This is okay, since the kmem code asks for pages after +dropping c_spinlock. The page_table_lock also nests with pagecache_lock and +pagemap_lru_lock spinlocks, and no code asks for memory with these locks +held. + +The page_table_lock is grabbed while holding the kernel_lock spinning monitor. + +The page_table_lock is a spin lock. + +Note: PTL can also be used to guarantee that no new clones using the +mm start up ... this is a loose form of stability on mm_users. For +example, it is used in copy_mm to protect against a racing tlb_gather_mmu +single address space optimization, so that the zap_page_range (from +truncate) does not lose sending ipi's to cloned threads that might +be spawned underneath it and go to user mode to drag in pte's into tlbs. + +swap_lock +-------------- +The swap devices are chained in priority order from the "swap_list" header. +The "swap_list" is used for the round-robin swaphandle allocation strategy. +The #free swaphandles is maintained in "nr_swap_pages". These two together +are protected by the swap_lock. + +The swap_lock also protects all the device reference counts on the +corresponding swaphandles, maintained in the "swap_map" array, and the +"highest_bit" and "lowest_bit" fields. + +The swap_lock is a spinlock, and is never acquired from intr level. + +To prevent races between swap space deletion or async readahead swapins +deciding whether a swap handle is being used, ie worthy of being read in +from disk, and an unmap -> swap_free making the handle unused, the swap +delete and readahead code grabs a temp reference on the swaphandle to +prevent warning messages from swap_duplicate <- read_swap_cache_async. + +Swap cache locking +------------------ +Pages are added into the swap cache with kernel_lock held, to make sure +that multiple pages are not being added (and hence lost) by associating +all of them with the same swaphandle. + +Pages are guaranteed not to be removed from the scache if the page is +"shared": ie, other processes hold reference on the page or the associated +swap handle. The only code that does not follow this rule is shrink_mmap, +which deletes pages from the swap cache if no process has a reference on +the page (multiple processes might have references on the corresponding +swap handle though). lookup_swap_cache() races with shrink_mmap, when +establishing a reference on a scache page, so, it must check whether the +page it located is still in the swapcache, or shrink_mmap deleted it. +(This race is due to the fact that shrink_mmap looks at the page ref +count with pagecache_lock, but then drops pagecache_lock before deleting +the page from the scache). + +do_wp_page and do_swap_page have MP races in them while trying to figure +out whether a page is "shared", by looking at the page_count + swap_count. +To preserve the sum of the counts, the page lock _must_ be acquired before +calling is_page_shared (else processes might switch their swap_count refs +to the page count refs, after the page count ref has been snapshotted). + +Swap device deletion code currently breaks all the scache assumptions, +since it grabs neither mmap_sem nor page_table_lock. diff --git a/Documentation/vm/numa b/Documentation/vm/numa new file mode 100644 index 00000000..ade01274 --- /dev/null +++ b/Documentation/vm/numa @@ -0,0 +1,149 @@ +Started Nov 1999 by Kanoj Sarcar <kanoj@sgi.com> + +What is NUMA? + +This question can be answered from a couple of perspectives: the +hardware view and the Linux software view. + +From the hardware perspective, a NUMA system is a computer platform that +comprises multiple components or assemblies each of which may contain 0 +or more CPUs, local memory, and/or IO buses. For brevity and to +disambiguate the hardware view of these physical components/assemblies +from the software abstraction thereof, we'll call the components/assemblies +'cells' in this document. + +Each of the 'cells' may be viewed as an SMP [symmetric multi-processor] subset +of the system--although some components necessary for a stand-alone SMP system +may not be populated on any given cell. The cells of the NUMA system are +connected together with some sort of system interconnect--e.g., a crossbar or +point-to-point link are common types of NUMA system interconnects. Both of +these types of interconnects can be aggregated to create NUMA platforms with +cells at multiple distances from other cells. + +For Linux, the NUMA platforms of interest are primarily what is known as Cache +Coherent NUMA or ccNUMA systems. With ccNUMA systems, all memory is visible +to and accessible from any CPU attached to any cell and cache coherency +is handled in hardware by the processor caches and/or the system interconnect. + +Memory access time and effective memory bandwidth varies depending on how far +away the cell containing the CPU or IO bus making the memory access is from the +cell containing the target memory. For example, access to memory by CPUs +attached to the same cell will experience faster access times and higher +bandwidths than accesses to memory on other, remote cells. NUMA platforms +can have cells at multiple remote distances from any given cell. + +Platform vendors don't build NUMA systems just to make software developers' +lives interesting. Rather, this architecture is a means to provide scalable +memory bandwidth. However, to achieve scalable memory bandwidth, system and +application software must arrange for a large majority of the memory references +[cache misses] to be to "local" memory--memory on the same cell, if any--or +to the closest cell with memory. + +This leads to the Linux software view of a NUMA system: + +Linux divides the system's hardware resources into multiple software +abstractions called "nodes". Linux maps the nodes onto the physical cells +of the hardware platform, abstracting away some of the details for some +architectures. As with physical cells, software nodes may contain 0 or more +CPUs, memory and/or IO buses. And, again, memory accesses to memory on +"closer" nodes--nodes that map to closer cells--will generally experience +faster access times and higher effective bandwidth than accesses to more +remote cells. + +For some architectures, such as x86, Linux will "hide" any node representing a +physical cell that has no memory attached, and reassign any CPUs attached to +that cell to a node representing a cell that does have memory. Thus, on +these architectures, one cannot assume that all CPUs that Linux associates with +a given node will see the same local memory access times and bandwidth. + +In addition, for some architectures, again x86 is an example, Linux supports +the emulation of additional nodes. For NUMA emulation, linux will carve up +the existing nodes--or the system memory for non-NUMA platforms--into multiple +nodes. Each emulated node will manage a fraction of the underlying cells' +physical memory. NUMA emluation is useful for testing NUMA kernel and +application features on non-NUMA platforms, and as a sort of memory resource +management mechanism when used together with cpusets. +[see Documentation/cgroups/cpusets.txt] + +For each node with memory, Linux constructs an independent memory management +subsystem, complete with its own free page lists, in-use page lists, usage +statistics and locks to mediate access. In addition, Linux constructs for +each memory zone [one or more of DMA, DMA32, NORMAL, HIGH_MEMORY, MOVABLE], +an ordered "zonelist". A zonelist specifies the zones/nodes to visit when a +selected zone/node cannot satisfy the allocation request. This situation, +when a zone has no available memory to satisfy a request, is called +"overflow" or "fallback". + +Because some nodes contain multiple zones containing different types of +memory, Linux must decide whether to order the zonelists such that allocations +fall back to the same zone type on a different node, or to a different zone +type on the same node. This is an important consideration because some zones, +such as DMA or DMA32, represent relatively scarce resources. Linux chooses +a default zonelist order based on the sizes of the various zone types relative +to the total memory of the node and the total memory of the system. The +default zonelist order may be overridden using the numa_zonelist_order kernel +boot parameter or sysctl. [see Documentation/kernel-parameters.txt and +Documentation/sysctl/vm.txt] + +By default, Linux will attempt to satisfy memory allocation requests from the +node to which the CPU that executes the request is assigned. Specifically, +Linux will attempt to allocate from the first node in the appropriate zonelist +for the node where the request originates. This is called "local allocation." +If the "local" node cannot satisfy the request, the kernel will examine other +nodes' zones in the selected zonelist looking for the first zone in the list +that can satisfy the request. + +Local allocation will tend to keep subsequent access to the allocated memory +"local" to the underlying physical resources and off the system interconnect-- +as long as the task on whose behalf the kernel allocated some memory does not +later migrate away from that memory. The Linux scheduler is aware of the +NUMA topology of the platform--embodied in the "scheduling domains" data +structures [see Documentation/scheduler/sched-domains.txt]--and the scheduler +attempts to minimize task migration to distant scheduling domains. However, +the scheduler does not take a task's NUMA footprint into account directly. +Thus, under sufficient imbalance, tasks can migrate between nodes, remote +from their initial node and kernel data structures. + +System administrators and application designers can restrict a task's migration +to improve NUMA locality using various CPU affinity command line interfaces, +such as taskset(1) and numactl(1), and program interfaces such as +sched_setaffinity(2). Further, one can modify the kernel's default local +allocation behavior using Linux NUMA memory policy. +[see Documentation/vm/numa_memory_policy.txt.] + +System administrators can restrict the CPUs and nodes' memories that a non- +privileged user can specify in the scheduling or NUMA commands and functions +using control groups and CPUsets. [see Documentation/cgroups/cpusets.txt] + +On architectures that do not hide memoryless nodes, Linux will include only +zones [nodes] with memory in the zonelists. This means that for a memoryless +node the "local memory node"--the node of the first zone in CPU's node's +zonelist--will not be the node itself. Rather, it will be the node that the +kernel selected as the nearest node with memory when it built the zonelists. +So, default, local allocations will succeed with the kernel supplying the +closest available memory. This is a consequence of the same mechanism that +allows such allocations to fallback to other nearby nodes when a node that +does contain memory overflows. + +Some kernel allocations do not want or cannot tolerate this allocation fallback +behavior. Rather they want to be sure they get memory from the specified node +or get notified that the node has no free memory. This is usually the case when +a subsystem allocates per CPU memory resources, for example. + +A typical model for making such an allocation is to obtain the node id of the +node to which the "current CPU" is attached using one of the kernel's +numa_node_id() or CPU_to_node() functions and then request memory from only +the node id returned. When such an allocation fails, the requesting subsystem +may revert to its own fallback path. The slab kernel memory allocator is an +example of this. Or, the subsystem may choose to disable or not to enable +itself on allocation failure. The kernel profiling subsystem is an example of +this. + +If the architecture supports--does not hide--memoryless nodes, then CPUs +attached to memoryless nodes would always incur the fallback path overhead +or some subsystems would fail to initialize if they attempted to allocated +memory exclusively from a node without memory. To support such +architectures transparently, kernel subsystems can use the numa_mem_id() +or cpu_to_mem() function to locate the "local memory node" for the calling or +specified CPU. Again, this is the same node from which default, local page +allocations will be attempted. diff --git a/Documentation/vm/numa_memory_policy.txt b/Documentation/vm/numa_memory_policy.txt new file mode 100644 index 00000000..4e7da654 --- /dev/null +++ b/Documentation/vm/numa_memory_policy.txt @@ -0,0 +1,453 @@ + +What is Linux Memory Policy? + +In the Linux kernel, "memory policy" determines from which node the kernel will +allocate memory in a NUMA system or in an emulated NUMA system. Linux has +supported platforms with Non-Uniform Memory Access architectures since 2.4.?. +The current memory policy support was added to Linux 2.6 around May 2004. This +document attempts to describe the concepts and APIs of the 2.6 memory policy +support. + +Memory policies should not be confused with cpusets +(Documentation/cgroups/cpusets.txt) +which is an administrative mechanism for restricting the nodes from which +memory may be allocated by a set of processes. Memory policies are a +programming interface that a NUMA-aware application can take advantage of. When +both cpusets and policies are applied to a task, the restrictions of the cpuset +takes priority. See "MEMORY POLICIES AND CPUSETS" below for more details. + +MEMORY POLICY CONCEPTS + +Scope of Memory Policies + +The Linux kernel supports _scopes_ of memory policy, described here from +most general to most specific: + + System Default Policy: this policy is "hard coded" into the kernel. It + is the policy that governs all page allocations that aren't controlled + by one of the more specific policy scopes discussed below. When the + system is "up and running", the system default policy will use "local + allocation" described below. However, during boot up, the system + default policy will be set to interleave allocations across all nodes + with "sufficient" memory, so as not to overload the initial boot node + with boot-time allocations. + + Task/Process Policy: this is an optional, per-task policy. When defined + for a specific task, this policy controls all page allocations made by or + on behalf of the task that aren't controlled by a more specific scope. + If a task does not define a task policy, then all page allocations that + would have been controlled by the task policy "fall back" to the System + Default Policy. + + The task policy applies to the entire address space of a task. Thus, + it is inheritable, and indeed is inherited, across both fork() + [clone() w/o the CLONE_VM flag] and exec*(). This allows a parent task + to establish the task policy for a child task exec()'d from an + executable image that has no awareness of memory policy. See the + MEMORY POLICY APIS section, below, for an overview of the system call + that a task may use to set/change its task/process policy. + + In a multi-threaded task, task policies apply only to the thread + [Linux kernel task] that installs the policy and any threads + subsequently created by that thread. Any sibling threads existing + at the time a new task policy is installed retain their current + policy. + + A task policy applies only to pages allocated after the policy is + installed. Any pages already faulted in by the task when the task + changes its task policy remain where they were allocated based on + the policy at the time they were allocated. + + VMA Policy: A "VMA" or "Virtual Memory Area" refers to a range of a task's + virtual address space. A task may define a specific policy for a range + of its virtual address space. See the MEMORY POLICIES APIS section, + below, for an overview of the mbind() system call used to set a VMA + policy. + + A VMA policy will govern the allocation of pages that back this region of + the address space. Any regions of the task's address space that don't + have an explicit VMA policy will fall back to the task policy, which may + itself fall back to the System Default Policy. + + VMA policies have a few complicating details: + + VMA policy applies ONLY to anonymous pages. These include pages + allocated for anonymous segments, such as the task stack and heap, and + any regions of the address space mmap()ed with the MAP_ANONYMOUS flag. + If a VMA policy is applied to a file mapping, it will be ignored if + the mapping used the MAP_SHARED flag. If the file mapping used the + MAP_PRIVATE flag, the VMA policy will only be applied when an + anonymous page is allocated on an attempt to write to the mapping-- + i.e., at Copy-On-Write. + + VMA policies are shared between all tasks that share a virtual address + space--a.k.a. threads--independent of when the policy is installed; and + they are inherited across fork(). However, because VMA policies refer + to a specific region of a task's address space, and because the address + space is discarded and recreated on exec*(), VMA policies are NOT + inheritable across exec(). Thus, only NUMA-aware applications may + use VMA policies. + + A task may install a new VMA policy on a sub-range of a previously + mmap()ed region. When this happens, Linux splits the existing virtual + memory area into 2 or 3 VMAs, each with it's own policy. + + By default, VMA policy applies only to pages allocated after the policy + is installed. Any pages already faulted into the VMA range remain + where they were allocated based on the policy at the time they were + allocated. However, since 2.6.16, Linux supports page migration via + the mbind() system call, so that page contents can be moved to match + a newly installed policy. + + Shared Policy: Conceptually, shared policies apply to "memory objects" + mapped shared into one or more tasks' distinct address spaces. An + application installs a shared policies the same way as VMA policies--using + the mbind() system call specifying a range of virtual addresses that map + the shared object. However, unlike VMA policies, which can be considered + to be an attribute of a range of a task's address space, shared policies + apply directly to the shared object. Thus, all tasks that attach to the + object share the policy, and all pages allocated for the shared object, + by any task, will obey the shared policy. + + As of 2.6.22, only shared memory segments, created by shmget() or + mmap(MAP_ANONYMOUS|MAP_SHARED), support shared policy. When shared + policy support was added to Linux, the associated data structures were + added to hugetlbfs shmem segments. At the time, hugetlbfs did not + support allocation at fault time--a.k.a lazy allocation--so hugetlbfs + shmem segments were never "hooked up" to the shared policy support. + Although hugetlbfs segments now support lazy allocation, their support + for shared policy has not been completed. + + As mentioned above [re: VMA policies], allocations of page cache + pages for regular files mmap()ed with MAP_SHARED ignore any VMA + policy installed on the virtual address range backed by the shared + file mapping. Rather, shared page cache pages, including pages backing + private mappings that have not yet been written by the task, follow + task policy, if any, else System Default Policy. + + The shared policy infrastructure supports different policies on subset + ranges of the shared object. However, Linux still splits the VMA of + the task that installs the policy for each range of distinct policy. + Thus, different tasks that attach to a shared memory segment can have + different VMA configurations mapping that one shared object. This + can be seen by examining the /proc/<pid>/numa_maps of tasks sharing + a shared memory region, when one task has installed shared policy on + one or more ranges of the region. + +Components of Memory Policies + + A Linux memory policy consists of a "mode", optional mode flags, and an + optional set of nodes. The mode determines the behavior of the policy, + the optional mode flags determine the behavior of the mode, and the + optional set of nodes can be viewed as the arguments to the policy + behavior. + + Internally, memory policies are implemented by a reference counted + structure, struct mempolicy. Details of this structure will be discussed + in context, below, as required to explain the behavior. + + Linux memory policy supports the following 4 behavioral modes: + + Default Mode--MPOL_DEFAULT: This mode is only used in the memory + policy APIs. Internally, MPOL_DEFAULT is converted to the NULL + memory policy in all policy scopes. Any existing non-default policy + will simply be removed when MPOL_DEFAULT is specified. As a result, + MPOL_DEFAULT means "fall back to the next most specific policy scope." + + For example, a NULL or default task policy will fall back to the + system default policy. A NULL or default vma policy will fall + back to the task policy. + + When specified in one of the memory policy APIs, the Default mode + does not use the optional set of nodes. + + It is an error for the set of nodes specified for this policy to + be non-empty. + + MPOL_BIND: This mode specifies that memory must come from the + set of nodes specified by the policy. Memory will be allocated from + the node in the set with sufficient free memory that is closest to + the node where the allocation takes place. + + MPOL_PREFERRED: This mode specifies that the allocation should be + attempted from the single node specified in the policy. If that + allocation fails, the kernel will search other nodes, in order of + increasing distance from the preferred node based on information + provided by the platform firmware. + containing the cpu where the allocation takes place. + + Internally, the Preferred policy uses a single node--the + preferred_node member of struct mempolicy. When the internal + mode flag MPOL_F_LOCAL is set, the preferred_node is ignored and + the policy is interpreted as local allocation. "Local" allocation + policy can be viewed as a Preferred policy that starts at the node + containing the cpu where the allocation takes place. + + It is possible for the user to specify that local allocation is + always preferred by passing an empty nodemask with this mode. + If an empty nodemask is passed, the policy cannot use the + MPOL_F_STATIC_NODES or MPOL_F_RELATIVE_NODES flags described + below. + + MPOL_INTERLEAVED: This mode specifies that page allocations be + interleaved, on a page granularity, across the nodes specified in + the policy. This mode also behaves slightly differently, based on + the context where it is used: + + For allocation of anonymous pages and shared memory pages, + Interleave mode indexes the set of nodes specified by the policy + using the page offset of the faulting address into the segment + [VMA] containing the address modulo the number of nodes specified + by the policy. It then attempts to allocate a page, starting at + the selected node, as if the node had been specified by a Preferred + policy or had been selected by a local allocation. That is, + allocation will follow the per node zonelist. + + For allocation of page cache pages, Interleave mode indexes the set + of nodes specified by the policy using a node counter maintained + per task. This counter wraps around to the lowest specified node + after it reaches the highest specified node. This will tend to + spread the pages out over the nodes specified by the policy based + on the order in which they are allocated, rather than based on any + page offset into an address range or file. During system boot up, + the temporary interleaved system default policy works in this + mode. + + Linux memory policy supports the following optional mode flags: + + MPOL_F_STATIC_NODES: This flag specifies that the nodemask passed by + the user should not be remapped if the task or VMA's set of allowed + nodes changes after the memory policy has been defined. + + Without this flag, anytime a mempolicy is rebound because of a + change in the set of allowed nodes, the node (Preferred) or + nodemask (Bind, Interleave) is remapped to the new set of + allowed nodes. This may result in nodes being used that were + previously undesired. + + With this flag, if the user-specified nodes overlap with the + nodes allowed by the task's cpuset, then the memory policy is + applied to their intersection. If the two sets of nodes do not + overlap, the Default policy is used. + + For example, consider a task that is attached to a cpuset with + mems 1-3 that sets an Interleave policy over the same set. If + the cpuset's mems change to 3-5, the Interleave will now occur + over nodes 3, 4, and 5. With this flag, however, since only node + 3 is allowed from the user's nodemask, the "interleave" only + occurs over that node. If no nodes from the user's nodemask are + now allowed, the Default behavior is used. + + MPOL_F_STATIC_NODES cannot be combined with the + MPOL_F_RELATIVE_NODES flag. It also cannot be used for + MPOL_PREFERRED policies that were created with an empty nodemask + (local allocation). + + MPOL_F_RELATIVE_NODES: This flag specifies that the nodemask passed + by the user will be mapped relative to the set of the task or VMA's + set of allowed nodes. The kernel stores the user-passed nodemask, + and if the allowed nodes changes, then that original nodemask will + be remapped relative to the new set of allowed nodes. + + Without this flag (and without MPOL_F_STATIC_NODES), anytime a + mempolicy is rebound because of a change in the set of allowed + nodes, the node (Preferred) or nodemask (Bind, Interleave) is + remapped to the new set of allowed nodes. That remap may not + preserve the relative nature of the user's passed nodemask to its + set of allowed nodes upon successive rebinds: a nodemask of + 1,3,5 may be remapped to 7-9 and then to 1-3 if the set of + allowed nodes is restored to its original state. + + With this flag, the remap is done so that the node numbers from + the user's passed nodemask are relative to the set of allowed + nodes. In other words, if nodes 0, 2, and 4 are set in the user's + nodemask, the policy will be effected over the first (and in the + Bind or Interleave case, the third and fifth) nodes in the set of + allowed nodes. The nodemask passed by the user represents nodes + relative to task or VMA's set of allowed nodes. + + If the user's nodemask includes nodes that are outside the range + of the new set of allowed nodes (for example, node 5 is set in + the user's nodemask when the set of allowed nodes is only 0-3), + then the remap wraps around to the beginning of the nodemask and, + if not already set, sets the node in the mempolicy nodemask. + + For example, consider a task that is attached to a cpuset with + mems 2-5 that sets an Interleave policy over the same set with + MPOL_F_RELATIVE_NODES. If the cpuset's mems change to 3-7, the + interleave now occurs over nodes 3,5-6. If the cpuset's mems + then change to 0,2-3,5, then the interleave occurs over nodes + 0,3,5. + + Thanks to the consistent remapping, applications preparing + nodemasks to specify memory policies using this flag should + disregard their current, actual cpuset imposed memory placement + and prepare the nodemask as if they were always located on + memory nodes 0 to N-1, where N is the number of memory nodes the + policy is intended to manage. Let the kernel then remap to the + set of memory nodes allowed by the task's cpuset, as that may + change over time. + + MPOL_F_RELATIVE_NODES cannot be combined with the + MPOL_F_STATIC_NODES flag. It also cannot be used for + MPOL_PREFERRED policies that were created with an empty nodemask + (local allocation). + +MEMORY POLICY REFERENCE COUNTING + +To resolve use/free races, struct mempolicy contains an atomic reference +count field. Internal interfaces, mpol_get()/mpol_put() increment and +decrement this reference count, respectively. mpol_put() will only free +the structure back to the mempolicy kmem cache when the reference count +goes to zero. + +When a new memory policy is allocated, its reference count is initialized +to '1', representing the reference held by the task that is installing the +new policy. When a pointer to a memory policy structure is stored in another +structure, another reference is added, as the task's reference will be dropped +on completion of the policy installation. + +During run-time "usage" of the policy, we attempt to minimize atomic operations +on the reference count, as this can lead to cache lines bouncing between cpus +and NUMA nodes. "Usage" here means one of the following: + +1) querying of the policy, either by the task itself [using the get_mempolicy() + API discussed below] or by another task using the /proc/<pid>/numa_maps + interface. + +2) examination of the policy to determine the policy mode and associated node + or node lists, if any, for page allocation. This is considered a "hot + path". Note that for MPOL_BIND, the "usage" extends across the entire + allocation process, which may sleep during page reclaimation, because the + BIND policy nodemask is used, by reference, to filter ineligible nodes. + +We can avoid taking an extra reference during the usages listed above as +follows: + +1) we never need to get/free the system default policy as this is never + changed nor freed, once the system is up and running. + +2) for querying the policy, we do not need to take an extra reference on the + target task's task policy nor vma policies because we always acquire the + task's mm's mmap_sem for read during the query. The set_mempolicy() and + mbind() APIs [see below] always acquire the mmap_sem for write when + installing or replacing task or vma policies. Thus, there is no possibility + of a task or thread freeing a policy while another task or thread is + querying it. + +3) Page allocation usage of task or vma policy occurs in the fault path where + we hold them mmap_sem for read. Again, because replacing the task or vma + policy requires that the mmap_sem be held for write, the policy can't be + freed out from under us while we're using it for page allocation. + +4) Shared policies require special consideration. One task can replace a + shared memory policy while another task, with a distinct mmap_sem, is + querying or allocating a page based on the policy. To resolve this + potential race, the shared policy infrastructure adds an extra reference + to the shared policy during lookup while holding a spin lock on the shared + policy management structure. This requires that we drop this extra + reference when we're finished "using" the policy. We must drop the + extra reference on shared policies in the same query/allocation paths + used for non-shared policies. For this reason, shared policies are marked + as such, and the extra reference is dropped "conditionally"--i.e., only + for shared policies. + + Because of this extra reference counting, and because we must lookup + shared policies in a tree structure under spinlock, shared policies are + more expensive to use in the page allocation path. This is especially + true for shared policies on shared memory regions shared by tasks running + on different NUMA nodes. This extra overhead can be avoided by always + falling back to task or system default policy for shared memory regions, + or by prefaulting the entire shared memory region into memory and locking + it down. However, this might not be appropriate for all applications. + +MEMORY POLICY APIs + +Linux supports 3 system calls for controlling memory policy. These APIS +always affect only the calling task, the calling task's address space, or +some shared object mapped into the calling task's address space. + + Note: the headers that define these APIs and the parameter data types + for user space applications reside in a package that is not part of + the Linux kernel. The kernel system call interfaces, with the 'sys_' + prefix, are defined in <linux/syscalls.h>; the mode and flag + definitions are defined in <linux/mempolicy.h>. + +Set [Task] Memory Policy: + + long set_mempolicy(int mode, const unsigned long *nmask, + unsigned long maxnode); + + Set's the calling task's "task/process memory policy" to mode + specified by the 'mode' argument and the set of nodes defined + by 'nmask'. 'nmask' points to a bit mask of node ids containing + at least 'maxnode' ids. Optional mode flags may be passed by + combining the 'mode' argument with the flag (for example: + MPOL_INTERLEAVE | MPOL_F_STATIC_NODES). + + See the set_mempolicy(2) man page for more details + + +Get [Task] Memory Policy or Related Information + + long get_mempolicy(int *mode, + const unsigned long *nmask, unsigned long maxnode, + void *addr, int flags); + + Queries the "task/process memory policy" of the calling task, or + the policy or location of a specified virtual address, depending + on the 'flags' argument. + + See the get_mempolicy(2) man page for more details + + +Install VMA/Shared Policy for a Range of Task's Address Space + + long mbind(void *start, unsigned long len, int mode, + const unsigned long *nmask, unsigned long maxnode, + unsigned flags); + + mbind() installs the policy specified by (mode, nmask, maxnodes) as + a VMA policy for the range of the calling task's address space + specified by the 'start' and 'len' arguments. Additional actions + may be requested via the 'flags' argument. + + See the mbind(2) man page for more details. + +MEMORY POLICY COMMAND LINE INTERFACE + +Although not strictly part of the Linux implementation of memory policy, +a command line tool, numactl(8), exists that allows one to: + ++ set the task policy for a specified program via set_mempolicy(2), fork(2) and + exec(2) + ++ set the shared policy for a shared memory segment via mbind(2) + +The numactl(8) tool is packaged with the run-time version of the library +containing the memory policy system call wrappers. Some distributions +package the headers and compile-time libraries in a separate development +package. + + +MEMORY POLICIES AND CPUSETS + +Memory policies work within cpusets as described above. For memory policies +that require a node or set of nodes, the nodes are restricted to the set of +nodes whose memories are allowed by the cpuset constraints. If the nodemask +specified for the policy contains nodes that are not allowed by the cpuset and +MPOL_F_RELATIVE_NODES is not used, the intersection of the set of nodes +specified for the policy and the set of nodes with memory is used. If the +result is the empty set, the policy is considered invalid and cannot be +installed. If MPOL_F_RELATIVE_NODES is used, the policy's nodes are mapped +onto and folded into the task's set of allowed nodes as previously described. + +The interaction of memory policies and cpusets can be problematic when tasks +in two cpusets share access to a memory region, such as shared memory segments +created by shmget() of mmap() with the MAP_ANONYMOUS and MAP_SHARED flags, and +any of the tasks install shared policy on the region, only nodes whose +memories are allowed in both cpusets may be used in the policies. Obtaining +this information requires "stepping outside" the memory policy APIs to use the +cpuset information and requires that one know in what cpusets other task might +be attaching to the shared region. Furthermore, if the cpusets' allowed +memory sets are disjoint, "local" allocation is the only valid policy. diff --git a/Documentation/vm/overcommit-accounting b/Documentation/vm/overcommit-accounting new file mode 100644 index 00000000..706d7ed9 --- /dev/null +++ b/Documentation/vm/overcommit-accounting @@ -0,0 +1,73 @@ +The Linux kernel supports the following overcommit handling modes + +0 - Heuristic overcommit handling. Obvious overcommits of + address space are refused. Used for a typical system. It + ensures a seriously wild allocation fails while allowing + overcommit to reduce swap usage. root is allowed to + allocate slightly more memory in this mode. This is the + default. + +1 - Always overcommit. Appropriate for some scientific + applications. + +2 - Don't overcommit. The total address space commit + for the system is not permitted to exceed swap + a + configurable percentage (default is 50) of physical RAM. + Depending on the percentage you use, in most situations + this means a process will not be killed while accessing + pages but will receive errors on memory allocation as + appropriate. + +The overcommit policy is set via the sysctl `vm.overcommit_memory'. + +The overcommit percentage is set via `vm.overcommit_ratio'. + +The current overcommit limit and amount committed are viewable in +/proc/meminfo as CommitLimit and Committed_AS respectively. + +Gotchas +------- + +The C language stack growth does an implicit mremap. If you want absolute +guarantees and run close to the edge you MUST mmap your stack for the +largest size you think you will need. For typical stack usage this does +not matter much but it's a corner case if you really really care + +In mode 2 the MAP_NORESERVE flag is ignored. + + +How It Works +------------ + +The overcommit is based on the following rules + +For a file backed map + SHARED or READ-only - 0 cost (the file is the map not swap) + PRIVATE WRITABLE - size of mapping per instance + +For an anonymous or /dev/zero map + SHARED - size of mapping + PRIVATE READ-only - 0 cost (but of little use) + PRIVATE WRITABLE - size of mapping per instance + +Additional accounting + Pages made writable copies by mmap + shmfs memory drawn from the same pool + +Status +------ + +o We account mmap memory mappings +o We account mprotect changes in commit +o We account mremap changes in size +o We account brk +o We account munmap +o We report the commit status in /proc +o Account and check on fork +o Review stack handling/building on exec +o SHMfs accounting +o Implement actual limit enforcement + +To Do +----- +o Account ptrace pages (this is hard) diff --git a/Documentation/vm/page_migration b/Documentation/vm/page_migration new file mode 100644 index 00000000..6513fe2d --- /dev/null +++ b/Documentation/vm/page_migration @@ -0,0 +1,149 @@ +Page migration +-------------- + +Page migration allows the moving of the physical location of pages between +nodes in a numa system while the process is running. This means that the +virtual addresses that the process sees do not change. However, the +system rearranges the physical location of those pages. + +The main intend of page migration is to reduce the latency of memory access +by moving pages near to the processor where the process accessing that memory +is running. + +Page migration allows a process to manually relocate the node on which its +pages are located through the MF_MOVE and MF_MOVE_ALL options while setting +a new memory policy via mbind(). The pages of process can also be relocated +from another process using the sys_migrate_pages() function call. The +migrate_pages function call takes two sets of nodes and moves pages of a +process that are located on the from nodes to the destination nodes. +Page migration functions are provided by the numactl package by Andi Kleen +(a version later than 0.9.3 is required. Get it from +ftp://oss.sgi.com/www/projects/libnuma/download/). numactl provides libnuma +which provides an interface similar to other numa functionality for page +migration. cat /proc/<pid>/numa_maps allows an easy review of where the +pages of a process are located. See also the numa_maps documentation in the +proc(5) man page. + +Manual migration is useful if for example the scheduler has relocated +a process to a processor on a distant node. A batch scheduler or an +administrator may detect the situation and move the pages of the process +nearer to the new processor. The kernel itself does only provide +manual page migration support. Automatic page migration may be implemented +through user space processes that move pages. A special function call +"move_pages" allows the moving of individual pages within a process. +A NUMA profiler may f.e. obtain a log showing frequent off node +accesses and may use the result to move pages to more advantageous +locations. + +Larger installations usually partition the system using cpusets into +sections of nodes. Paul Jackson has equipped cpusets with the ability to +move pages when a task is moved to another cpuset (See +Documentation/cgroups/cpusets.txt). +Cpusets allows the automation of process locality. If a task is moved to +a new cpuset then also all its pages are moved with it so that the +performance of the process does not sink dramatically. Also the pages +of processes in a cpuset are moved if the allowed memory nodes of a +cpuset are changed. + +Page migration allows the preservation of the relative location of pages +within a group of nodes for all migration techniques which will preserve a +particular memory allocation pattern generated even after migrating a +process. This is necessary in order to preserve the memory latencies. +Processes will run with similar performance after migration. + +Page migration occurs in several steps. First a high level +description for those trying to use migrate_pages() from the kernel +(for userspace usage see the Andi Kleen's numactl package mentioned above) +and then a low level description of how the low level details work. + +A. In kernel use of migrate_pages() +----------------------------------- + +1. Remove pages from the LRU. + + Lists of pages to be migrated are generated by scanning over + pages and moving them into lists. This is done by + calling isolate_lru_page(). + Calling isolate_lru_page increases the references to the page + so that it cannot vanish while the page migration occurs. + It also prevents the swapper or other scans to encounter + the page. + +2. We need to have a function of type new_page_t that can be + passed to migrate_pages(). This function should figure out + how to allocate the correct new page given the old page. + +3. The migrate_pages() function is called which attempts + to do the migration. It will call the function to allocate + the new page for each page that is considered for + moving. + +B. How migrate_pages() works +---------------------------- + +migrate_pages() does several passes over its list of pages. A page is moved +if all references to a page are removable at the time. The page has +already been removed from the LRU via isolate_lru_page() and the refcount +is increased so that the page cannot be freed while page migration occurs. + +Steps: + +1. Lock the page to be migrated + +2. Insure that writeback is complete. + +3. Prep the new page that we want to move to. It is locked + and set to not being uptodate so that all accesses to the new + page immediately lock while the move is in progress. + +4. The new page is prepped with some settings from the old page so that + accesses to the new page will discover a page with the correct settings. + +5. All the page table references to the page are converted + to migration entries or dropped (nonlinear vmas). + This decrease the mapcount of a page. If the resulting + mapcount is not zero then we do not migrate the page. + All user space processes that attempt to access the page + will now wait on the page lock. + +6. The radix tree lock is taken. This will cause all processes trying + to access the page via the mapping to block on the radix tree spinlock. + +7. The refcount of the page is examined and we back out if references remain + otherwise we know that we are the only one referencing this page. + +8. The radix tree is checked and if it does not contain the pointer to this + page then we back out because someone else modified the radix tree. + +9. The radix tree is changed to point to the new page. + +10. The reference count of the old page is dropped because the radix tree + reference is gone. A reference to the new page is established because + the new page is referenced to by the radix tree. + +11. The radix tree lock is dropped. With that lookups in the mapping + become possible again. Processes will move from spinning on the tree_lock + to sleeping on the locked new page. + +12. The page contents are copied to the new page. + +13. The remaining page flags are copied to the new page. + +14. The old page flags are cleared to indicate that the page does + not provide any information anymore. + +15. Queued up writeback on the new page is triggered. + +16. If migration entries were page then replace them with real ptes. Doing + so will enable access for user space processes not already waiting for + the page lock. + +19. The page locks are dropped from the old and new page. + Processes waiting on the page lock will redo their page faults + and will reach the new page. + +20. The new page is moved to the LRU and can be scanned by the swapper + etc again. + +Christoph Lameter, May 8, 2006. + diff --git a/Documentation/vm/pagemap.txt b/Documentation/vm/pagemap.txt new file mode 100644 index 00000000..4600cbe3 --- /dev/null +++ b/Documentation/vm/pagemap.txt @@ -0,0 +1,151 @@ +pagemap, from the userspace perspective +--------------------------------------- + +pagemap is a new (as of 2.6.25) set of interfaces in the kernel that allow +userspace programs to examine the page tables and related information by +reading files in /proc. + +There are three components to pagemap: + + * /proc/pid/pagemap. This file lets a userspace process find out which + physical frame each virtual page is mapped to. It contains one 64-bit + value for each virtual page, containing the following data (from + fs/proc/task_mmu.c, above pagemap_read): + + * Bits 0-54 page frame number (PFN) if present + * Bits 0-4 swap type if swapped + * Bits 5-54 swap offset if swapped + * Bits 55-60 page shift (page size = 1<<page shift) + * Bit 61 reserved for future use + * Bit 62 page swapped + * Bit 63 page present + + If the page is not present but in swap, then the PFN contains an + encoding of the swap file number and the page's offset into the + swap. Unmapped pages return a null PFN. This allows determining + precisely which pages are mapped (or in swap) and comparing mapped + pages between processes. + + Efficient users of this interface will use /proc/pid/maps to + determine which areas of memory are actually mapped and llseek to + skip over unmapped regions. + + * /proc/kpagecount. This file contains a 64-bit count of the number of + times each page is mapped, indexed by PFN. + + * /proc/kpageflags. This file contains a 64-bit set of flags for each + page, indexed by PFN. + + The flags are (from fs/proc/page.c, above kpageflags_read): + + 0. LOCKED + 1. ERROR + 2. REFERENCED + 3. UPTODATE + 4. DIRTY + 5. LRU + 6. ACTIVE + 7. SLAB + 8. WRITEBACK + 9. RECLAIM + 10. BUDDY + 11. MMAP + 12. ANON + 13. SWAPCACHE + 14. SWAPBACKED + 15. COMPOUND_HEAD + 16. COMPOUND_TAIL + 16. HUGE + 18. UNEVICTABLE + 19. HWPOISON + 20. NOPAGE + 21. KSM + 22. THP + +Short descriptions to the page flags: + + 0. LOCKED + page is being locked for exclusive access, eg. by undergoing read/write IO + + 7. SLAB + page is managed by the SLAB/SLOB/SLUB/SLQB kernel memory allocator + When compound page is used, SLUB/SLQB will only set this flag on the head + page; SLOB will not flag it at all. + +10. BUDDY + a free memory block managed by the buddy system allocator + The buddy system organizes free memory in blocks of various orders. + An order N block has 2^N physically contiguous pages, with the BUDDY flag + set for and _only_ for the first page. + +15. COMPOUND_HEAD +16. COMPOUND_TAIL + A compound page with order N consists of 2^N physically contiguous pages. + A compound page with order 2 takes the form of "HTTT", where H donates its + head page and T donates its tail page(s). The major consumers of compound + pages are hugeTLB pages (Documentation/vm/hugetlbpage.txt), the SLUB etc. + memory allocators and various device drivers. However in this interface, + only huge/giga pages are made visible to end users. +17. HUGE + this is an integral part of a HugeTLB page + +19. HWPOISON + hardware detected memory corruption on this page: don't touch the data! + +20. NOPAGE + no page frame exists at the requested address + +21. KSM + identical memory pages dynamically shared between one or more processes + +22. THP + contiguous pages which construct transparent hugepages + + [IO related page flags] + 1. ERROR IO error occurred + 3. UPTODATE page has up-to-date data + ie. for file backed page: (in-memory data revision >= on-disk one) + 4. DIRTY page has been written to, hence contains new data + ie. for file backed page: (in-memory data revision > on-disk one) + 8. WRITEBACK page is being synced to disk + + [LRU related page flags] + 5. LRU page is in one of the LRU lists + 6. ACTIVE page is in the active LRU list +18. UNEVICTABLE page is in the unevictable (non-)LRU list + It is somehow pinned and not a candidate for LRU page reclaims, + eg. ramfs pages, shmctl(SHM_LOCK) and mlock() memory segments + 2. REFERENCED page has been referenced since last LRU list enqueue/requeue + 9. RECLAIM page will be reclaimed soon after its pageout IO completed +11. MMAP a memory mapped page +12. ANON a memory mapped page that is not part of a file +13. SWAPCACHE page is mapped to swap space, ie. has an associated swap entry +14. SWAPBACKED page is backed by swap/RAM + +The page-types tool in this directory can be used to query the above flags. + +Using pagemap to do something useful: + +The general procedure for using pagemap to find out about a process' memory +usage goes like this: + + 1. Read /proc/pid/maps to determine which parts of the memory space are + mapped to what. + 2. Select the maps you are interested in -- all of them, or a particular + library, or the stack or the heap, etc. + 3. Open /proc/pid/pagemap and seek to the pages you would like to examine. + 4. Read a u64 for each page from pagemap. + 5. Open /proc/kpagecount and/or /proc/kpageflags. For each PFN you just + read, seek to that entry in the file, and read the data you want. + +For example, to find the "unique set size" (USS), which is the amount of +memory that a process is using that is not shared with any other process, +you can go through every map in the process, find the PFNs, look those up +in kpagecount, and tally up the number of pages that are only referenced +once. + +Other notes: + +Reading from any of the files will return -EINVAL if you are not starting +the read on an 8-byte boundary (e.g., if you seeked an odd number of bytes +into the file), or if the size of the read is not a multiple of 8 bytes. diff --git a/Documentation/vm/slub.txt b/Documentation/vm/slub.txt new file mode 100644 index 00000000..6752870c --- /dev/null +++ b/Documentation/vm/slub.txt @@ -0,0 +1,283 @@ +Short users guide for SLUB +-------------------------- + +The basic philosophy of SLUB is very different from SLAB. SLAB +requires rebuilding the kernel to activate debug options for all +slab caches. SLUB always includes full debugging but it is off by default. +SLUB can enable debugging only for selected slabs in order to avoid +an impact on overall system performance which may make a bug more +difficult to find. + +In order to switch debugging on one can add a option "slub_debug" +to the kernel command line. That will enable full debugging for +all slabs. + +Typically one would then use the "slabinfo" command to get statistical +data and perform operation on the slabs. By default slabinfo only lists +slabs that have data in them. See "slabinfo -h" for more options when +running the command. slabinfo can be compiled with + +gcc -o slabinfo tools/slub/slabinfo.c + +Some of the modes of operation of slabinfo require that slub debugging +be enabled on the command line. F.e. no tracking information will be +available without debugging on and validation can only partially +be performed if debugging was not switched on. + +Some more sophisticated uses of slub_debug: +------------------------------------------- + +Parameters may be given to slub_debug. If none is specified then full +debugging is enabled. Format: + +slub_debug=<Debug-Options> Enable options for all slabs +slub_debug=<Debug-Options>,<slab name> + Enable options only for select slabs + +Possible debug options are + F Sanity checks on (enables SLAB_DEBUG_FREE. Sorry + SLAB legacy issues) + Z Red zoning + P Poisoning (object and padding) + U User tracking (free and alloc) + T Trace (please only use on single slabs) + A Toggle failslab filter mark for the cache + O Switch debugging off for caches that would have + caused higher minimum slab orders + - Switch all debugging off (useful if the kernel is + configured with CONFIG_SLUB_DEBUG_ON) + +F.e. in order to boot just with sanity checks and red zoning one would specify: + + slub_debug=FZ + +Trying to find an issue in the dentry cache? Try + + slub_debug=,dentry + +to only enable debugging on the dentry cache. + +Red zoning and tracking may realign the slab. We can just apply sanity checks +to the dentry cache with + + slub_debug=F,dentry + +Debugging options may require the minimum possible slab order to increase as +a result of storing the metadata (for example, caches with PAGE_SIZE object +sizes). This has a higher liklihood of resulting in slab allocation errors +in low memory situations or if there's high fragmentation of memory. To +switch off debugging for such caches by default, use + + slub_debug=O + +In case you forgot to enable debugging on the kernel command line: It is +possible to enable debugging manually when the kernel is up. Look at the +contents of: + +/sys/kernel/slab/<slab name>/ + +Look at the writable files. Writing 1 to them will enable the +corresponding debug option. All options can be set on a slab that does +not contain objects. If the slab already contains objects then sanity checks +and tracing may only be enabled. The other options may cause the realignment +of objects. + +Careful with tracing: It may spew out lots of information and never stop if +used on the wrong slab. + +Slab merging +------------ + +If no debug options are specified then SLUB may merge similar slabs together +in order to reduce overhead and increase cache hotness of objects. +slabinfo -a displays which slabs were merged together. + +Slab validation +--------------- + +SLUB can validate all object if the kernel was booted with slub_debug. In +order to do so you must have the slabinfo tool. Then you can do + +slabinfo -v + +which will test all objects. Output will be generated to the syslog. + +This also works in a more limited way if boot was without slab debug. +In that case slabinfo -v simply tests all reachable objects. Usually +these are in the cpu slabs and the partial slabs. Full slabs are not +tracked by SLUB in a non debug situation. + +Getting more performance +------------------------ + +To some degree SLUB's performance is limited by the need to take the +list_lock once in a while to deal with partial slabs. That overhead is +governed by the order of the allocation for each slab. The allocations +can be influenced by kernel parameters: + +slub_min_objects=x (default 4) +slub_min_order=x (default 0) +slub_max_order=x (default 3 (PAGE_ALLOC_COSTLY_ORDER)) + +slub_min_objects allows to specify how many objects must at least fit +into one slab in order for the allocation order to be acceptable. +In general slub will be able to perform this number of allocations +on a slab without consulting centralized resources (list_lock) where +contention may occur. + +slub_min_order specifies a minim order of slabs. A similar effect like +slub_min_objects. + +slub_max_order specified the order at which slub_min_objects should no +longer be checked. This is useful to avoid SLUB trying to generate +super large order pages to fit slub_min_objects of a slab cache with +large object sizes into one high order page. Setting command line +parameter debug_guardpage_minorder=N (N > 0), forces setting +slub_max_order to 0, what cause minimum possible order of slabs +allocation. + +SLUB Debug output +----------------- + +Here is a sample of slub debug output: + +==================================================================== +BUG kmalloc-8: Redzone overwritten +-------------------------------------------------------------------- + +INFO: 0xc90f6d28-0xc90f6d2b. First byte 0x00 instead of 0xcc +INFO: Slab 0xc528c530 flags=0x400000c3 inuse=61 fp=0xc90f6d58 +INFO: Object 0xc90f6d20 @offset=3360 fp=0xc90f6d58 +INFO: Allocated in get_modalias+0x61/0xf5 age=53 cpu=1 pid=554 + +Bytes b4 0xc90f6d10: 00 00 00 00 00 00 00 00 5a 5a 5a 5a 5a 5a 5a 5a ........ZZZZZZZZ + Object 0xc90f6d20: 31 30 31 39 2e 30 30 35 1019.005 + Redzone 0xc90f6d28: 00 cc cc cc . + Padding 0xc90f6d50: 5a 5a 5a 5a 5a 5a 5a 5a ZZZZZZZZ + + [<c010523d>] dump_trace+0x63/0x1eb + [<c01053df>] show_trace_log_lvl+0x1a/0x2f + [<c010601d>] show_trace+0x12/0x14 + [<c0106035>] dump_stack+0x16/0x18 + [<c017e0fa>] object_err+0x143/0x14b + [<c017e2cc>] check_object+0x66/0x234 + [<c017eb43>] __slab_free+0x239/0x384 + [<c017f446>] kfree+0xa6/0xc6 + [<c02e2335>] get_modalias+0xb9/0xf5 + [<c02e23b7>] dmi_dev_uevent+0x27/0x3c + [<c027866a>] dev_uevent+0x1ad/0x1da + [<c0205024>] kobject_uevent_env+0x20a/0x45b + [<c020527f>] kobject_uevent+0xa/0xf + [<c02779f1>] store_uevent+0x4f/0x58 + [<c027758e>] dev_attr_store+0x29/0x2f + [<c01bec4f>] sysfs_write_file+0x16e/0x19c + [<c0183ba7>] vfs_write+0xd1/0x15a + [<c01841d7>] sys_write+0x3d/0x72 + [<c0104112>] sysenter_past_esp+0x5f/0x99 + [<b7f7b410>] 0xb7f7b410 + ======================= + +FIX kmalloc-8: Restoring Redzone 0xc90f6d28-0xc90f6d2b=0xcc + +If SLUB encounters a corrupted object (full detection requires the kernel +to be booted with slub_debug) then the following output will be dumped +into the syslog: + +1. Description of the problem encountered + +This will be a message in the system log starting with + +=============================================== +BUG <slab cache affected>: <What went wrong> +----------------------------------------------- + +INFO: <corruption start>-<corruption_end> <more info> +INFO: Slab <address> <slab information> +INFO: Object <address> <object information> +INFO: Allocated in <kernel function> age=<jiffies since alloc> cpu=<allocated by + cpu> pid=<pid of the process> +INFO: Freed in <kernel function> age=<jiffies since free> cpu=<freed by cpu> + pid=<pid of the process> + +(Object allocation / free information is only available if SLAB_STORE_USER is +set for the slab. slub_debug sets that option) + +2. The object contents if an object was involved. + +Various types of lines can follow the BUG SLUB line: + +Bytes b4 <address> : <bytes> + Shows a few bytes before the object where the problem was detected. + Can be useful if the corruption does not stop with the start of the + object. + +Object <address> : <bytes> + The bytes of the object. If the object is inactive then the bytes + typically contain poison values. Any non-poison value shows a + corruption by a write after free. + +Redzone <address> : <bytes> + The Redzone following the object. The Redzone is used to detect + writes after the object. All bytes should always have the same + value. If there is any deviation then it is due to a write after + the object boundary. + + (Redzone information is only available if SLAB_RED_ZONE is set. + slub_debug sets that option) + +Padding <address> : <bytes> + Unused data to fill up the space in order to get the next object + properly aligned. In the debug case we make sure that there are + at least 4 bytes of padding. This allows the detection of writes + before the object. + +3. A stackdump + +The stackdump describes the location where the error was detected. The cause +of the corruption is may be more likely found by looking at the function that +allocated or freed the object. + +4. Report on how the problem was dealt with in order to ensure the continued +operation of the system. + +These are messages in the system log beginning with + +FIX <slab cache affected>: <corrective action taken> + +In the above sample SLUB found that the Redzone of an active object has +been overwritten. Here a string of 8 characters was written into a slab that +has the length of 8 characters. However, a 8 character string needs a +terminating 0. That zero has overwritten the first byte of the Redzone field. +After reporting the details of the issue encountered the FIX SLUB message +tells us that SLUB has restored the Redzone to its proper value and then +system operations continue. + +Emergency operations: +--------------------- + +Minimal debugging (sanity checks alone) can be enabled by booting with + + slub_debug=F + +This will be generally be enough to enable the resiliency features of slub +which will keep the system running even if a bad kernel component will +keep corrupting objects. This may be important for production systems. +Performance will be impacted by the sanity checks and there will be a +continual stream of error messages to the syslog but no additional memory +will be used (unlike full debugging). + +No guarantees. The kernel component still needs to be fixed. Performance +may be optimized further by locating the slab that experiences corruption +and enabling debugging only for that cache + +I.e. + + slub_debug=F,dentry + +If the corruption occurs by writing after the end of the object then it +may be advisable to enable a Redzone to avoid corrupting the beginning +of other objects. + + slub_debug=FZ,dentry + +Christoph Lameter, May 30, 2007 diff --git a/Documentation/vm/transhuge.txt b/Documentation/vm/transhuge.txt new file mode 100644 index 00000000..29bdf62a --- /dev/null +++ b/Documentation/vm/transhuge.txt @@ -0,0 +1,299 @@ += Transparent Hugepage Support = + +== Objective == + +Performance critical computing applications dealing with large memory +working sets are already running on top of libhugetlbfs and in turn +hugetlbfs. Transparent Hugepage Support is an alternative means of +using huge pages for the backing of virtual memory with huge pages +that supports the automatic promotion and demotion of page sizes and +without the shortcomings of hugetlbfs. + +Currently it only works for anonymous memory mappings but in the +future it can expand over the pagecache layer starting with tmpfs. + +The reason applications are running faster is because of two +factors. The first factor is almost completely irrelevant and it's not +of significant interest because it'll also have the downside of +requiring larger clear-page copy-page in page faults which is a +potentially negative effect. The first factor consists in taking a +single page fault for each 2M virtual region touched by userland (so +reducing the enter/exit kernel frequency by a 512 times factor). This +only matters the first time the memory is accessed for the lifetime of +a memory mapping. The second long lasting and much more important +factor will affect all subsequent accesses to the memory for the whole +runtime of the application. The second factor consist of two +components: 1) the TLB miss will run faster (especially with +virtualization using nested pagetables but almost always also on bare +metal without virtualization) and 2) a single TLB entry will be +mapping a much larger amount of virtual memory in turn reducing the +number of TLB misses. With virtualization and nested pagetables the +TLB can be mapped of larger size only if both KVM and the Linux guest +are using hugepages but a significant speedup already happens if only +one of the two is using hugepages just because of the fact the TLB +miss is going to run faster. + +== Design == + +- "graceful fallback": mm components which don't have transparent + hugepage knowledge fall back to breaking a transparent hugepage and + working on the regular pages and their respective regular pmd/pte + mappings + +- if a hugepage allocation fails because of memory fragmentation, + regular pages should be gracefully allocated instead and mixed in + the same vma without any failure or significant delay and without + userland noticing + +- if some task quits and more hugepages become available (either + immediately in the buddy or through the VM), guest physical memory + backed by regular pages should be relocated on hugepages + automatically (with khugepaged) + +- it doesn't require memory reservation and in turn it uses hugepages + whenever possible (the only possible reservation here is kernelcore= + to avoid unmovable pages to fragment all the memory but such a tweak + is not specific to transparent hugepage support and it's a generic + feature that applies to all dynamic high order allocations in the + kernel) + +- this initial support only offers the feature in the anonymous memory + regions but it'd be ideal to move it to tmpfs and the pagecache + later + +Transparent Hugepage Support maximizes the usefulness of free memory +if compared to the reservation approach of hugetlbfs by allowing all +unused memory to be used as cache or other movable (or even unmovable +entities). It doesn't require reservation to prevent hugepage +allocation failures to be noticeable from userland. It allows paging +and all other advanced VM features to be available on the +hugepages. It requires no modifications for applications to take +advantage of it. + +Applications however can be further optimized to take advantage of +this feature, like for example they've been optimized before to avoid +a flood of mmap system calls for every malloc(4k). Optimizing userland +is by far not mandatory and khugepaged already can take care of long +lived page allocations even for hugepage unaware applications that +deals with large amounts of memory. + +In certain cases when hugepages are enabled system wide, application +may end up allocating more memory resources. An application may mmap a +large region but only touch 1 byte of it, in that case a 2M page might +be allocated instead of a 4k page for no good. This is why it's +possible to disable hugepages system-wide and to only have them inside +MADV_HUGEPAGE madvise regions. + +Embedded systems should enable hugepages only inside madvise regions +to eliminate any risk of wasting any precious byte of memory and to +only run faster. + +Applications that gets a lot of benefit from hugepages and that don't +risk to lose memory by using hugepages, should use +madvise(MADV_HUGEPAGE) on their critical mmapped regions. + +== sysfs == + +Transparent Hugepage Support can be entirely disabled (mostly for +debugging purposes) or only enabled inside MADV_HUGEPAGE regions (to +avoid the risk of consuming more memory resources) or enabled system +wide. This can be achieved with one of: + +echo always >/sys/kernel/mm/transparent_hugepage/enabled +echo madvise >/sys/kernel/mm/transparent_hugepage/enabled +echo never >/sys/kernel/mm/transparent_hugepage/enabled + +It's also possible to limit defrag efforts in the VM to generate +hugepages in case they're not immediately free to madvise regions or +to never try to defrag memory and simply fallback to regular pages +unless hugepages are immediately available. Clearly if we spend CPU +time to defrag memory, we would expect to gain even more by the fact +we use hugepages later instead of regular pages. This isn't always +guaranteed, but it may be more likely in case the allocation is for a +MADV_HUGEPAGE region. + +echo always >/sys/kernel/mm/transparent_hugepage/defrag +echo madvise >/sys/kernel/mm/transparent_hugepage/defrag +echo never >/sys/kernel/mm/transparent_hugepage/defrag + +khugepaged will be automatically started when +transparent_hugepage/enabled is set to "always" or "madvise, and it'll +be automatically shutdown if it's set to "never". + +khugepaged runs usually at low frequency so while one may not want to +invoke defrag algorithms synchronously during the page faults, it +should be worth invoking defrag at least in khugepaged. However it's +also possible to disable defrag in khugepaged by writing 0 or enable +defrag in khugepaged by writing 1: + +echo 0 >/sys/kernel/mm/transparent_hugepage/khugepaged/defrag +echo 1 >/sys/kernel/mm/transparent_hugepage/khugepaged/defrag + +You can also control how many pages khugepaged should scan at each +pass: + +/sys/kernel/mm/transparent_hugepage/khugepaged/pages_to_scan + +and how many milliseconds to wait in khugepaged between each pass (you +can set this to 0 to run khugepaged at 100% utilization of one core): + +/sys/kernel/mm/transparent_hugepage/khugepaged/scan_sleep_millisecs + +and how many milliseconds to wait in khugepaged if there's an hugepage +allocation failure to throttle the next allocation attempt. + +/sys/kernel/mm/transparent_hugepage/khugepaged/alloc_sleep_millisecs + +The khugepaged progress can be seen in the number of pages collapsed: + +/sys/kernel/mm/transparent_hugepage/khugepaged/pages_collapsed + +for each pass: + +/sys/kernel/mm/transparent_hugepage/khugepaged/full_scans + +== Boot parameter == + +You can change the sysfs boot time defaults of Transparent Hugepage +Support by passing the parameter "transparent_hugepage=always" or +"transparent_hugepage=madvise" or "transparent_hugepage=never" +(without "") to the kernel command line. + +== Need of application restart == + +The transparent_hugepage/enabled values only affect future +behavior. So to make them effective you need to restart any +application that could have been using hugepages. This also applies to +the regions registered in khugepaged. + +== get_user_pages and follow_page == + +get_user_pages and follow_page if run on a hugepage, will return the +head or tail pages as usual (exactly as they would do on +hugetlbfs). Most gup users will only care about the actual physical +address of the page and its temporary pinning to release after the I/O +is complete, so they won't ever notice the fact the page is huge. But +if any driver is going to mangle over the page structure of the tail +page (like for checking page->mapping or other bits that are relevant +for the head page and not the tail page), it should be updated to jump +to check head page instead (while serializing properly against +split_huge_page() to avoid the head and tail pages to disappear from +under it, see the futex code to see an example of that, hugetlbfs also +needed special handling in futex code for similar reasons). + +NOTE: these aren't new constraints to the GUP API, and they match the +same constrains that applies to hugetlbfs too, so any driver capable +of handling GUP on hugetlbfs will also work fine on transparent +hugepage backed mappings. + +In case you can't handle compound pages if they're returned by +follow_page, the FOLL_SPLIT bit can be specified as parameter to +follow_page, so that it will split the hugepages before returning +them. Migration for example passes FOLL_SPLIT as parameter to +follow_page because it's not hugepage aware and in fact it can't work +at all on hugetlbfs (but it instead works fine on transparent +hugepages thanks to FOLL_SPLIT). migration simply can't deal with +hugepages being returned (as it's not only checking the pfn of the +page and pinning it during the copy but it pretends to migrate the +memory in regular page sizes and with regular pte/pmd mappings). + +== Optimizing the applications == + +To be guaranteed that the kernel will map a 2M page immediately in any +memory region, the mmap region has to be hugepage naturally +aligned. posix_memalign() can provide that guarantee. + +== Hugetlbfs == + +You can use hugetlbfs on a kernel that has transparent hugepage +support enabled just fine as always. No difference can be noted in +hugetlbfs other than there will be less overall fragmentation. All +usual features belonging to hugetlbfs are preserved and +unaffected. libhugetlbfs will also work fine as usual. + +== Graceful fallback == + +Code walking pagetables but unware about huge pmds can simply call +split_huge_page_pmd(mm, pmd) where the pmd is the one returned by +pmd_offset. It's trivial to make the code transparent hugepage aware +by just grepping for "pmd_offset" and adding split_huge_page_pmd where +missing after pmd_offset returns the pmd. Thanks to the graceful +fallback design, with a one liner change, you can avoid to write +hundred if not thousand of lines of complex code to make your code +hugepage aware. + +If you're not walking pagetables but you run into a physical hugepage +but you can't handle it natively in your code, you can split it by +calling split_huge_page(page). This is what the Linux VM does before +it tries to swapout the hugepage for example. + +Example to make mremap.c transparent hugepage aware with a one liner +change: + +diff --git a/mm/mremap.c b/mm/mremap.c +--- a/mm/mremap.c ++++ b/mm/mremap.c +@@ -41,6 +41,7 @@ static pmd_t *get_old_pmd(struct mm_stru + return NULL; + + pmd = pmd_offset(pud, addr); ++ split_huge_page_pmd(mm, pmd); + if (pmd_none_or_clear_bad(pmd)) + return NULL; + +== Locking in hugepage aware code == + +We want as much code as possible hugepage aware, as calling +split_huge_page() or split_huge_page_pmd() has a cost. + +To make pagetable walks huge pmd aware, all you need to do is to call +pmd_trans_huge() on the pmd returned by pmd_offset. You must hold the +mmap_sem in read (or write) mode to be sure an huge pmd cannot be +created from under you by khugepaged (khugepaged collapse_huge_page +takes the mmap_sem in write mode in addition to the anon_vma lock). If +pmd_trans_huge returns false, you just fallback in the old code +paths. If instead pmd_trans_huge returns true, you have to take the +mm->page_table_lock and re-run pmd_trans_huge. Taking the +page_table_lock will prevent the huge pmd to be converted into a +regular pmd from under you (split_huge_page can run in parallel to the +pagetable walk). If the second pmd_trans_huge returns false, you +should just drop the page_table_lock and fallback to the old code as +before. Otherwise you should run pmd_trans_splitting on the pmd. In +case pmd_trans_splitting returns true, it means split_huge_page is +already in the middle of splitting the page. So if pmd_trans_splitting +returns true it's enough to drop the page_table_lock and call +wait_split_huge_page and then fallback the old code paths. You are +guaranteed by the time wait_split_huge_page returns, the pmd isn't +huge anymore. If pmd_trans_splitting returns false, you can proceed to +process the huge pmd and the hugepage natively. Once finished you can +drop the page_table_lock. + +== compound_lock, get_user_pages and put_page == + +split_huge_page internally has to distribute the refcounts in the head +page to the tail pages before clearing all PG_head/tail bits from the +page structures. It can do that easily for refcounts taken by huge pmd +mappings. But the GUI API as created by hugetlbfs (that returns head +and tail pages if running get_user_pages on an address backed by any +hugepage), requires the refcount to be accounted on the tail pages and +not only in the head pages, if we want to be able to run +split_huge_page while there are gup pins established on any tail +page. Failure to be able to run split_huge_page if there's any gup pin +on any tail page, would mean having to split all hugepages upfront in +get_user_pages which is unacceptable as too many gup users are +performance critical and they must work natively on hugepages like +they work natively on hugetlbfs already (hugetlbfs is simpler because +hugetlbfs pages cannot be splitted so there wouldn't be requirement of +accounting the pins on the tail pages for hugetlbfs). If we wouldn't +account the gup refcounts on the tail pages during gup, we won't know +anymore which tail page is pinned by gup and which is not while we run +split_huge_page. But we still have to add the gup pin to the head page +too, to know when we can free the compound page in case it's never +splitted during its lifetime. That requires changing not just +get_page, but put_page as well so that when put_page runs on a tail +page (and only on a tail page) it will find its respective head page, +and then it will decrease the head page refcount in addition to the +tail page refcount. To obtain a head page reliably and to decrease its +refcount without race conditions, put_page has to serialize against +__split_huge_page_refcount using a special per-page lock called +compound_lock. diff --git a/Documentation/vm/unevictable-lru.txt b/Documentation/vm/unevictable-lru.txt new file mode 100644 index 00000000..fa206ccc --- /dev/null +++ b/Documentation/vm/unevictable-lru.txt @@ -0,0 +1,690 @@ + ============================== + UNEVICTABLE LRU INFRASTRUCTURE + ============================== + +======== +CONTENTS +======== + + (*) The Unevictable LRU + + - The unevictable page list. + - Memory control group interaction. + - Marking address spaces unevictable. + - Detecting Unevictable Pages. + - vmscan's handling of unevictable pages. + + (*) mlock()'d pages. + + - History. + - Basic management. + - mlock()/mlockall() system call handling. + - Filtering special vmas. + - munlock()/munlockall() system call handling. + - Migrating mlocked pages. + - mmap(MAP_LOCKED) system call handling. + - munmap()/exit()/exec() system call handling. + - try_to_unmap(). + - try_to_munlock() reverse map scan. + - Page reclaim in shrink_*_list(). + + +============ +INTRODUCTION +============ + +This document describes the Linux memory manager's "Unevictable LRU" +infrastructure and the use of this to manage several types of "unevictable" +pages. + +The document attempts to provide the overall rationale behind this mechanism +and the rationale for some of the design decisions that drove the +implementation. The latter design rationale is discussed in the context of an +implementation description. Admittedly, one can obtain the implementation +details - the "what does it do?" - by reading the code. One hopes that the +descriptions below add value by provide the answer to "why does it do that?". + + +=================== +THE UNEVICTABLE LRU +=================== + +The Unevictable LRU facility adds an additional LRU list to track unevictable +pages and to hide these pages from vmscan. This mechanism is based on a patch +by Larry Woodman of Red Hat to address several scalability problems with page +reclaim in Linux. The problems have been observed at customer sites on large +memory x86_64 systems. + +To illustrate this with an example, a non-NUMA x86_64 platform with 128GB of +main memory will have over 32 million 4k pages in a single zone. When a large +fraction of these pages are not evictable for any reason [see below], vmscan +will spend a lot of time scanning the LRU lists looking for the small fraction +of pages that are evictable. This can result in a situation where all CPUs are +spending 100% of their time in vmscan for hours or days on end, with the system +completely unresponsive. + +The unevictable list addresses the following classes of unevictable pages: + + (*) Those owned by ramfs. + + (*) Those mapped into SHM_LOCK'd shared memory regions. + + (*) Those mapped into VM_LOCKED [mlock()ed] VMAs. + +The infrastructure may also be able to handle other conditions that make pages +unevictable, either by definition or by circumstance, in the future. + + +THE UNEVICTABLE PAGE LIST +------------------------- + +The Unevictable LRU infrastructure consists of an additional, per-zone, LRU list +called the "unevictable" list and an associated page flag, PG_unevictable, to +indicate that the page is being managed on the unevictable list. + +The PG_unevictable flag is analogous to, and mutually exclusive with, the +PG_active flag in that it indicates on which LRU list a page resides when +PG_lru is set. + +The Unevictable LRU infrastructure maintains unevictable pages on an additional +LRU list for a few reasons: + + (1) We get to "treat unevictable pages just like we treat other pages in the + system - which means we get to use the same code to manipulate them, the + same code to isolate them (for migrate, etc.), the same code to keep track + of the statistics, etc..." [Rik van Riel] + + (2) We want to be able to migrate unevictable pages between nodes for memory + defragmentation, workload management and memory hotplug. The linux kernel + can only migrate pages that it can successfully isolate from the LRU + lists. If we were to maintain pages elsewhere than on an LRU-like list, + where they can be found by isolate_lru_page(), we would prevent their + migration, unless we reworked migration code to find the unevictable pages + itself. + + +The unevictable list does not differentiate between file-backed and anonymous, +swap-backed pages. This differentiation is only important while the pages are, +in fact, evictable. + +The unevictable list benefits from the "arrayification" of the per-zone LRU +lists and statistics originally proposed and posted by Christoph Lameter. + +The unevictable list does not use the LRU pagevec mechanism. Rather, +unevictable pages are placed directly on the page's zone's unevictable list +under the zone lru_lock. This allows us to prevent the stranding of pages on +the unevictable list when one task has the page isolated from the LRU and other +tasks are changing the "evictability" state of the page. + + +MEMORY CONTROL GROUP INTERACTION +-------------------------------- + +The unevictable LRU facility interacts with the memory control group [aka +memory controller; see Documentation/cgroups/memory.txt] by extending the +lru_list enum. + +The memory controller data structure automatically gets a per-zone unevictable +list as a result of the "arrayification" of the per-zone LRU lists (one per +lru_list enum element). The memory controller tracks the movement of pages to +and from the unevictable list. + +When a memory control group comes under memory pressure, the controller will +not attempt to reclaim pages on the unevictable list. This has a couple of +effects: + + (1) Because the pages are "hidden" from reclaim on the unevictable list, the + reclaim process can be more efficient, dealing only with pages that have a + chance of being reclaimed. + + (2) On the other hand, if too many of the pages charged to the control group + are unevictable, the evictable portion of the working set of the tasks in + the control group may not fit into the available memory. This can cause + the control group to thrash or to OOM-kill tasks. + + +MARKING ADDRESS SPACES UNEVICTABLE +---------------------------------- + +For facilities such as ramfs none of the pages attached to the address space +may be evicted. To prevent eviction of any such pages, the AS_UNEVICTABLE +address space flag is provided, and this can be manipulated by a filesystem +using a number of wrapper functions: + + (*) void mapping_set_unevictable(struct address_space *mapping); + + Mark the address space as being completely unevictable. + + (*) void mapping_clear_unevictable(struct address_space *mapping); + + Mark the address space as being evictable. + + (*) int mapping_unevictable(struct address_space *mapping); + + Query the address space, and return true if it is completely + unevictable. + +These are currently used in two places in the kernel: + + (1) By ramfs to mark the address spaces of its inodes when they are created, + and this mark remains for the life of the inode. + + (2) By SYSV SHM to mark SHM_LOCK'd address spaces until SHM_UNLOCK is called. + + Note that SHM_LOCK is not required to page in the locked pages if they're + swapped out; the application must touch the pages manually if it wants to + ensure they're in memory. + + +DETECTING UNEVICTABLE PAGES +--------------------------- + +The function page_evictable() in vmscan.c determines whether a page is +evictable or not using the query function outlined above [see section "Marking +address spaces unevictable"] to check the AS_UNEVICTABLE flag. + +For address spaces that are so marked after being populated (as SHM regions +might be), the lock action (eg: SHM_LOCK) can be lazy, and need not populate +the page tables for the region as does, for example, mlock(), nor need it make +any special effort to push any pages in the SHM_LOCK'd area to the unevictable +list. Instead, vmscan will do this if and when it encounters the pages during +a reclamation scan. + +On an unlock action (such as SHM_UNLOCK), the unlocker (eg: shmctl()) must scan +the pages in the region and "rescue" them from the unevictable list if no other +condition is keeping them unevictable. If an unevictable region is destroyed, +the pages are also "rescued" from the unevictable list in the process of +freeing them. + +page_evictable() also checks for mlocked pages by testing an additional page +flag, PG_mlocked (as wrapped by PageMlocked()). If the page is NOT mlocked, +and a non-NULL VMA is supplied, page_evictable() will check whether the VMA is +VM_LOCKED via is_mlocked_vma(). is_mlocked_vma() will SetPageMlocked() and +update the appropriate statistics if the vma is VM_LOCKED. This method allows +efficient "culling" of pages in the fault path that are being faulted in to +VM_LOCKED VMAs. + + +VMSCAN'S HANDLING OF UNEVICTABLE PAGES +-------------------------------------- + +If unevictable pages are culled in the fault path, or moved to the unevictable +list at mlock() or mmap() time, vmscan will not encounter the pages until they +have become evictable again (via munlock() for example) and have been "rescued" +from the unevictable list. However, there may be situations where we decide, +for the sake of expediency, to leave a unevictable page on one of the regular +active/inactive LRU lists for vmscan to deal with. vmscan checks for such +pages in all of the shrink_{active|inactive|page}_list() functions and will +"cull" such pages that it encounters: that is, it diverts those pages to the +unevictable list for the zone being scanned. + +There may be situations where a page is mapped into a VM_LOCKED VMA, but the +page is not marked as PG_mlocked. Such pages will make it all the way to +shrink_page_list() where they will be detected when vmscan walks the reverse +map in try_to_unmap(). If try_to_unmap() returns SWAP_MLOCK, +shrink_page_list() will cull the page at that point. + +To "cull" an unevictable page, vmscan simply puts the page back on the LRU list +using putback_lru_page() - the inverse operation to isolate_lru_page() - after +dropping the page lock. Because the condition which makes the page unevictable +may change once the page is unlocked, putback_lru_page() will recheck the +unevictable state of a page that it places on the unevictable list. If the +page has become unevictable, putback_lru_page() removes it from the list and +retries, including the page_unevictable() test. Because such a race is a rare +event and movement of pages onto the unevictable list should be rare, these +extra evictabilty checks should not occur in the majority of calls to +putback_lru_page(). + + +============= +MLOCKED PAGES +============= + +The unevictable page list is also useful for mlock(), in addition to ramfs and +SYSV SHM. Note that mlock() is only available in CONFIG_MMU=y situations; in +NOMMU situations, all mappings are effectively mlocked. + + +HISTORY +------- + +The "Unevictable mlocked Pages" infrastructure is based on work originally +posted by Nick Piggin in an RFC patch entitled "mm: mlocked pages off LRU". +Nick posted his patch as an alternative to a patch posted by Christoph Lameter +to achieve the same objective: hiding mlocked pages from vmscan. + +In Nick's patch, he used one of the struct page LRU list link fields as a count +of VM_LOCKED VMAs that map the page. This use of the link field for a count +prevented the management of the pages on an LRU list, and thus mlocked pages +were not migratable as isolate_lru_page() could not find them, and the LRU list +link field was not available to the migration subsystem. + +Nick resolved this by putting mlocked pages back on the lru list before +attempting to isolate them, thus abandoning the count of VM_LOCKED VMAs. When +Nick's patch was integrated with the Unevictable LRU work, the count was +replaced by walking the reverse map to determine whether any VM_LOCKED VMAs +mapped the page. More on this below. + + +BASIC MANAGEMENT +---------------- + +mlocked pages - pages mapped into a VM_LOCKED VMA - are a class of unevictable +pages. When such a page has been "noticed" by the memory management subsystem, +the page is marked with the PG_mlocked flag. This can be manipulated using the +PageMlocked() functions. + +A PG_mlocked page will be placed on the unevictable list when it is added to +the LRU. Such pages can be "noticed" by memory management in several places: + + (1) in the mlock()/mlockall() system call handlers; + + (2) in the mmap() system call handler when mmapping a region with the + MAP_LOCKED flag; + + (3) mmapping a region in a task that has called mlockall() with the MCL_FUTURE + flag + + (4) in the fault path, if mlocked pages are "culled" in the fault path, + and when a VM_LOCKED stack segment is expanded; or + + (5) as mentioned above, in vmscan:shrink_page_list() when attempting to + reclaim a page in a VM_LOCKED VMA via try_to_unmap() + +all of which result in the VM_LOCKED flag being set for the VMA if it doesn't +already have it set. + +mlocked pages become unlocked and rescued from the unevictable list when: + + (1) mapped in a range unlocked via the munlock()/munlockall() system calls; + + (2) munmap()'d out of the last VM_LOCKED VMA that maps the page, including + unmapping at task exit; + + (3) when the page is truncated from the last VM_LOCKED VMA of an mmapped file; + or + + (4) before a page is COW'd in a VM_LOCKED VMA. + + +mlock()/mlockall() SYSTEM CALL HANDLING +--------------------------------------- + +Both [do_]mlock() and [do_]mlockall() system call handlers call mlock_fixup() +for each VMA in the range specified by the call. In the case of mlockall(), +this is the entire active address space of the task. Note that mlock_fixup() +is used for both mlocking and munlocking a range of memory. A call to mlock() +an already VM_LOCKED VMA, or to munlock() a VMA that is not VM_LOCKED is +treated as a no-op, and mlock_fixup() simply returns. + +If the VMA passes some filtering as described in "Filtering Special Vmas" +below, mlock_fixup() will attempt to merge the VMA with its neighbors or split +off a subset of the VMA if the range does not cover the entire VMA. Once the +VMA has been merged or split or neither, mlock_fixup() will call +__mlock_vma_pages_range() to fault in the pages via get_user_pages() and to +mark the pages as mlocked via mlock_vma_page(). + +Note that the VMA being mlocked might be mapped with PROT_NONE. In this case, +get_user_pages() will be unable to fault in the pages. That's okay. If pages +do end up getting faulted into this VM_LOCKED VMA, we'll handle them in the +fault path or in vmscan. + +Also note that a page returned by get_user_pages() could be truncated or +migrated out from under us, while we're trying to mlock it. To detect this, +__mlock_vma_pages_range() checks page_mapping() after acquiring the page lock. +If the page is still associated with its mapping, we'll go ahead and call +mlock_vma_page(). If the mapping is gone, we just unlock the page and move on. +In the worst case, this will result in a page mapped in a VM_LOCKED VMA +remaining on a normal LRU list without being PageMlocked(). Again, vmscan will +detect and cull such pages. + +mlock_vma_page() will call TestSetPageMlocked() for each page returned by +get_user_pages(). We use TestSetPageMlocked() because the page might already +be mlocked by another task/VMA and we don't want to do extra work. We +especially do not want to count an mlocked page more than once in the +statistics. If the page was already mlocked, mlock_vma_page() need do nothing +more. + +If the page was NOT already mlocked, mlock_vma_page() attempts to isolate the +page from the LRU, as it is likely on the appropriate active or inactive list +at that time. If the isolate_lru_page() succeeds, mlock_vma_page() will put +back the page - by calling putback_lru_page() - which will notice that the page +is now mlocked and divert the page to the zone's unevictable list. If +mlock_vma_page() is unable to isolate the page from the LRU, vmscan will handle +it later if and when it attempts to reclaim the page. + + +FILTERING SPECIAL VMAS +---------------------- + +mlock_fixup() filters several classes of "special" VMAs: + +1) VMAs with VM_IO or VM_PFNMAP set are skipped entirely. The pages behind + these mappings are inherently pinned, so we don't need to mark them as + mlocked. In any case, most of the pages have no struct page in which to so + mark the page. Because of this, get_user_pages() will fail for these VMAs, + so there is no sense in attempting to visit them. + +2) VMAs mapping hugetlbfs page are already effectively pinned into memory. We + neither need nor want to mlock() these pages. However, to preserve the + prior behavior of mlock() - before the unevictable/mlock changes - + mlock_fixup() will call make_pages_present() in the hugetlbfs VMA range to + allocate the huge pages and populate the ptes. + +3) VMAs with VM_DONTEXPAND or VM_RESERVED are generally userspace mappings of + kernel pages, such as the VDSO page, relay channel pages, etc. These pages + are inherently unevictable and are not managed on the LRU lists. + mlock_fixup() treats these VMAs the same as hugetlbfs VMAs. It calls + make_pages_present() to populate the ptes. + +Note that for all of these special VMAs, mlock_fixup() does not set the +VM_LOCKED flag. Therefore, we won't have to deal with them later during +munlock(), munmap() or task exit. Neither does mlock_fixup() account these +VMAs against the task's "locked_vm". + + +munlock()/munlockall() SYSTEM CALL HANDLING +------------------------------------------- + +The munlock() and munlockall() system calls are handled by the same functions - +do_mlock[all]() - as the mlock() and mlockall() system calls with the unlock vs +lock operation indicated by an argument. So, these system calls are also +handled by mlock_fixup(). Again, if called for an already munlocked VMA, +mlock_fixup() simply returns. Because of the VMA filtering discussed above, +VM_LOCKED will not be set in any "special" VMAs. So, these VMAs will be +ignored for munlock. + +If the VMA is VM_LOCKED, mlock_fixup() again attempts to merge or split off the +specified range. The range is then munlocked via the function +__mlock_vma_pages_range() - the same function used to mlock a VMA range - +passing a flag to indicate that munlock() is being performed. + +Because the VMA access protections could have been changed to PROT_NONE after +faulting in and mlocking pages, get_user_pages() was unreliable for visiting +these pages for munlocking. Because we don't want to leave pages mlocked, +get_user_pages() was enhanced to accept a flag to ignore the permissions when +fetching the pages - all of which should be resident as a result of previous +mlocking. + +For munlock(), __mlock_vma_pages_range() unlocks individual pages by calling +munlock_vma_page(). munlock_vma_page() unconditionally clears the PG_mlocked +flag using TestClearPageMlocked(). As with mlock_vma_page(), +munlock_vma_page() use the Test*PageMlocked() function to handle the case where +the page might have already been unlocked by another task. If the page was +mlocked, munlock_vma_page() updates that zone statistics for the number of +mlocked pages. Note, however, that at this point we haven't checked whether +the page is mapped by other VM_LOCKED VMAs. + +We can't call try_to_munlock(), the function that walks the reverse map to +check for other VM_LOCKED VMAs, without first isolating the page from the LRU. +try_to_munlock() is a variant of try_to_unmap() and thus requires that the page +not be on an LRU list [more on these below]. However, the call to +isolate_lru_page() could fail, in which case we couldn't try_to_munlock(). So, +we go ahead and clear PG_mlocked up front, as this might be the only chance we +have. If we can successfully isolate the page, we go ahead and +try_to_munlock(), which will restore the PG_mlocked flag and update the zone +page statistics if it finds another VMA holding the page mlocked. If we fail +to isolate the page, we'll have left a potentially mlocked page on the LRU. +This is fine, because we'll catch it later if and if vmscan tries to reclaim +the page. This should be relatively rare. + + +MIGRATING MLOCKED PAGES +----------------------- + +A page that is being migrated has been isolated from the LRU lists and is held +locked across unmapping of the page, updating the page's address space entry +and copying the contents and state, until the page table entry has been +replaced with an entry that refers to the new page. Linux supports migration +of mlocked pages and other unevictable pages. This involves simply moving the +PG_mlocked and PG_unevictable states from the old page to the new page. + +Note that page migration can race with mlocking or munlocking of the same page. +This has been discussed from the mlock/munlock perspective in the respective +sections above. Both processes (migration and m[un]locking) hold the page +locked. This provides the first level of synchronization. Page migration +zeros out the page_mapping of the old page before unlocking it, so m[un]lock +can skip these pages by testing the page mapping under page lock. + +To complete page migration, we place the new and old pages back onto the LRU +after dropping the page lock. The "unneeded" page - old page on success, new +page on failure - will be freed when the reference count held by the migration +process is released. To ensure that we don't strand pages on the unevictable +list because of a race between munlock and migration, page migration uses the +putback_lru_page() function to add migrated pages back to the LRU. + + +mmap(MAP_LOCKED) SYSTEM CALL HANDLING +------------------------------------- + +In addition the the mlock()/mlockall() system calls, an application can request +that a region of memory be mlocked supplying the MAP_LOCKED flag to the mmap() +call. Furthermore, any mmap() call or brk() call that expands the heap by a +task that has previously called mlockall() with the MCL_FUTURE flag will result +in the newly mapped memory being mlocked. Before the unevictable/mlock +changes, the kernel simply called make_pages_present() to allocate pages and +populate the page table. + +To mlock a range of memory under the unevictable/mlock infrastructure, the +mmap() handler and task address space expansion functions call +mlock_vma_pages_range() specifying the vma and the address range to mlock. +mlock_vma_pages_range() filters VMAs like mlock_fixup(), as described above in +"Filtering Special VMAs". It will clear the VM_LOCKED flag, which will have +already been set by the caller, in filtered VMAs. Thus these VMA's need not be +visited for munlock when the region is unmapped. + +For "normal" VMAs, mlock_vma_pages_range() calls __mlock_vma_pages_range() to +fault/allocate the pages and mlock them. Again, like mlock_fixup(), +mlock_vma_pages_range() downgrades the mmap semaphore to read mode before +attempting to fault/allocate and mlock the pages and "upgrades" the semaphore +back to write mode before returning. + +The callers of mlock_vma_pages_range() will have already added the memory range +to be mlocked to the task's "locked_vm". To account for filtered VMAs, +mlock_vma_pages_range() returns the number of pages NOT mlocked. All of the +callers then subtract a non-negative return value from the task's locked_vm. A +negative return value represent an error - for example, from get_user_pages() +attempting to fault in a VMA with PROT_NONE access. In this case, we leave the +memory range accounted as locked_vm, as the protections could be changed later +and pages allocated into that region. + + +munmap()/exit()/exec() SYSTEM CALL HANDLING +------------------------------------------- + +When unmapping an mlocked region of memory, whether by an explicit call to +munmap() or via an internal unmap from exit() or exec() processing, we must +munlock the pages if we're removing the last VM_LOCKED VMA that maps the pages. +Before the unevictable/mlock changes, mlocking did not mark the pages in any +way, so unmapping them required no processing. + +To munlock a range of memory under the unevictable/mlock infrastructure, the +munmap() handler and task address space call tear down function +munlock_vma_pages_all(). The name reflects the observation that one always +specifies the entire VMA range when munlock()ing during unmap of a region. +Because of the VMA filtering when mlocking() regions, only "normal" VMAs that +actually contain mlocked pages will be passed to munlock_vma_pages_all(). + +munlock_vma_pages_all() clears the VM_LOCKED VMA flag and, like mlock_fixup() +for the munlock case, calls __munlock_vma_pages_range() to walk the page table +for the VMA's memory range and munlock_vma_page() each resident page mapped by +the VMA. This effectively munlocks the page, only if this is the last +VM_LOCKED VMA that maps the page. + + +try_to_unmap() +-------------- + +Pages can, of course, be mapped into multiple VMAs. Some of these VMAs may +have VM_LOCKED flag set. It is possible for a page mapped into one or more +VM_LOCKED VMAs not to have the PG_mlocked flag set and therefore reside on one +of the active or inactive LRU lists. This could happen if, for example, a task +in the process of munlocking the page could not isolate the page from the LRU. +As a result, vmscan/shrink_page_list() might encounter such a page as described +in section "vmscan's handling of unevictable pages". To handle this situation, +try_to_unmap() checks for VM_LOCKED VMAs while it is walking a page's reverse +map. + +try_to_unmap() is always called, by either vmscan for reclaim or for page +migration, with the argument page locked and isolated from the LRU. Separate +functions handle anonymous and mapped file pages, as these types of pages have +different reverse map mechanisms. + + (*) try_to_unmap_anon() + + To unmap anonymous pages, each VMA in the list anchored in the anon_vma + must be visited - at least until a VM_LOCKED VMA is encountered. If the + page is being unmapped for migration, VM_LOCKED VMAs do not stop the + process because mlocked pages are migratable. However, for reclaim, if + the page is mapped into a VM_LOCKED VMA, the scan stops. + + try_to_unmap_anon() attempts to acquire in read mode the mmap semaphore of + the mm_struct to which the VMA belongs. If this is successful, it will + mlock the page via mlock_vma_page() - we wouldn't have gotten to + try_to_unmap_anon() if the page were already mlocked - and will return + SWAP_MLOCK, indicating that the page is unevictable. + + If the mmap semaphore cannot be acquired, we are not sure whether the page + is really unevictable or not. In this case, try_to_unmap_anon() will + return SWAP_AGAIN. + + (*) try_to_unmap_file() - linear mappings + + Unmapping of a mapped file page works the same as for anonymous mappings, + except that the scan visits all VMAs that map the page's index/page offset + in the page's mapping's reverse map priority search tree. It also visits + each VMA in the page's mapping's non-linear list, if the list is + non-empty. + + As for anonymous pages, on encountering a VM_LOCKED VMA for a mapped file + page, try_to_unmap_file() will attempt to acquire the associated + mm_struct's mmap semaphore to mlock the page, returning SWAP_MLOCK if this + is successful, and SWAP_AGAIN, if not. + + (*) try_to_unmap_file() - non-linear mappings + + If a page's mapping contains a non-empty non-linear mapping VMA list, then + try_to_un{map|lock}() must also visit each VMA in that list to determine + whether the page is mapped in a VM_LOCKED VMA. Again, the scan must visit + all VMAs in the non-linear list to ensure that the pages is not/should not + be mlocked. + + If a VM_LOCKED VMA is found in the list, the scan could terminate. + However, there is no easy way to determine whether the page is actually + mapped in a given VMA - either for unmapping or testing whether the + VM_LOCKED VMA actually pins the page. + + try_to_unmap_file() handles non-linear mappings by scanning a certain + number of pages - a "cluster" - in each non-linear VMA associated with the + page's mapping, for each file mapped page that vmscan tries to unmap. If + this happens to unmap the page we're trying to unmap, try_to_unmap() will + notice this on return (page_mapcount(page) will be 0) and return + SWAP_SUCCESS. Otherwise, it will return SWAP_AGAIN, causing vmscan to + recirculate this page. We take advantage of the cluster scan in + try_to_unmap_cluster() as follows: + + For each non-linear VMA, try_to_unmap_cluster() attempts to acquire the + mmap semaphore of the associated mm_struct for read without blocking. + + If this attempt is successful and the VMA is VM_LOCKED, + try_to_unmap_cluster() will retain the mmap semaphore for the scan; + otherwise it drops it here. + + Then, for each page in the cluster, if we're holding the mmap semaphore + for a locked VMA, try_to_unmap_cluster() calls mlock_vma_page() to + mlock the page. This call is a no-op if the page is already locked, + but will mlock any pages in the non-linear mapping that happen to be + unlocked. + + If one of the pages so mlocked is the page passed in to try_to_unmap(), + try_to_unmap_cluster() will return SWAP_MLOCK, rather than the default + SWAP_AGAIN. This will allow vmscan to cull the page, rather than + recirculating it on the inactive list. + + Again, if try_to_unmap_cluster() cannot acquire the VMA's mmap sem, it + returns SWAP_AGAIN, indicating that the page is mapped by a VM_LOCKED + VMA, but couldn't be mlocked. + + +try_to_munlock() REVERSE MAP SCAN +--------------------------------- + + [!] TODO/FIXME: a better name might be page_mlocked() - analogous to the + page_referenced() reverse map walker. + +When munlock_vma_page() [see section "munlock()/munlockall() System Call +Handling" above] tries to munlock a page, it needs to determine whether or not +the page is mapped by any VM_LOCKED VMA without actually attempting to unmap +all PTEs from the page. For this purpose, the unevictable/mlock infrastructure +introduced a variant of try_to_unmap() called try_to_munlock(). + +try_to_munlock() calls the same functions as try_to_unmap() for anonymous and +mapped file pages with an additional argument specifying unlock versus unmap +processing. Again, these functions walk the respective reverse maps looking +for VM_LOCKED VMAs. When such a VMA is found for anonymous pages and file +pages mapped in linear VMAs, as in the try_to_unmap() case, the functions +attempt to acquire the associated mmap semaphore, mlock the page via +mlock_vma_page() and return SWAP_MLOCK. This effectively undoes the +pre-clearing of the page's PG_mlocked done by munlock_vma_page. + +If try_to_unmap() is unable to acquire a VM_LOCKED VMA's associated mmap +semaphore, it will return SWAP_AGAIN. This will allow shrink_page_list() to +recycle the page on the inactive list and hope that it has better luck with the +page next time. + +For file pages mapped into non-linear VMAs, the try_to_munlock() logic works +slightly differently. On encountering a VM_LOCKED non-linear VMA that might +map the page, try_to_munlock() returns SWAP_AGAIN without actually mlocking the +page. munlock_vma_page() will just leave the page unlocked and let vmscan deal +with it - the usual fallback position. + +Note that try_to_munlock()'s reverse map walk must visit every VMA in a page's +reverse map to determine that a page is NOT mapped into any VM_LOCKED VMA. +However, the scan can terminate when it encounters a VM_LOCKED VMA and can +successfully acquire the VMA's mmap semaphore for read and mlock the page. +Although try_to_munlock() might be called a great many times when munlocking a +large region or tearing down a large address space that has been mlocked via +mlockall(), overall this is a fairly rare event. + + +PAGE RECLAIM IN shrink_*_list() +------------------------------- + +shrink_active_list() culls any obviously unevictable pages - i.e. +!page_evictable(page, NULL) - diverting these to the unevictable list. +However, shrink_active_list() only sees unevictable pages that made it onto the +active/inactive lru lists. Note that these pages do not have PageUnevictable +set - otherwise they would be on the unevictable list and shrink_active_list +would never see them. + +Some examples of these unevictable pages on the LRU lists are: + + (1) ramfs pages that have been placed on the LRU lists when first allocated. + + (2) SHM_LOCK'd shared memory pages. shmctl(SHM_LOCK) does not attempt to + allocate or fault in the pages in the shared memory region. This happens + when an application accesses the page the first time after SHM_LOCK'ing + the segment. + + (3) mlocked pages that could not be isolated from the LRU and moved to the + unevictable list in mlock_vma_page(). + + (4) Pages mapped into multiple VM_LOCKED VMAs, but try_to_munlock() couldn't + acquire the VMA's mmap semaphore to test the flags and set PageMlocked. + munlock_vma_page() was forced to let the page back on to the normal LRU + list for vmscan to handle. + +shrink_inactive_list() also diverts any unevictable pages that it finds on the +inactive lists to the appropriate zone's unevictable list. + +shrink_inactive_list() should only see SHM_LOCK'd pages that became SHM_LOCK'd +after shrink_active_list() had moved them to the inactive list, or pages mapped +into VM_LOCKED VMAs that munlock_vma_page() couldn't isolate from the LRU to +recheck via try_to_munlock(). shrink_inactive_list() won't notice the latter, +but will pass on to shrink_page_list(). + +shrink_page_list() again culls obviously unevictable pages that it could +encounter for similar reason to shrink_inactive_list(). Pages mapped into +VM_LOCKED VMAs but without PG_mlocked set will make it all the way to +try_to_unmap(). shrink_page_list() will divert them to the unevictable list +when try_to_unmap() returns SWAP_MLOCK, as discussed above. |